Method for locating sector data in a memory disk by examining a plurality of headers near an initial pointer

ABSTRACT

A method of allocating free physical memory in a solid state memory disk for a sector of data of a given size is described. Allocation begins by determining the amount of free physical memory in each block of the solid state memory disk. Afterward, the block with the greatest amount of free physical memory is selected. Sufficient free memory to store the sector of data is reserved in the selected block, provided the amount of free memory within selected block exceeds the given size of the sector of data. 
     A method of allocating free physical memory for sectors of user data stored within a block that is the focus of clean-up is also described. 
     A method of locating a sector of data stored in a solid state memory disk given a sector number and a pointer to a block sector translation table of a block is also described.

FIELD OF THE INVENTION

The present invention pertains to the field of computer storage systems.More particularly, the present invention pertains to a method andcircuitry for locating sectors of user data in a solid state memorydisk.

BACKGROUND OF THE INVENTION

Most prior personal computers include many types of memory storagedevices. Hard magnetic disk drives are used typically for mass storage,while different types of semiconductor memories are used for otherpurposes.

The non-volatility of hard magnetic disk drives is one of the chiefreasons for their use. They may be turned on and off without loss ofdata. Hard drives store data on concentric tracks. Each track includesseveral sectors, each of which is typically 512 bytes in length. Readsand writes to magnetic disk drives occur one bit at a time.

A typical user file stored on a magnetic disk drive occupies manysectors, randomly located on the surface of the disk drive. A fileallocation table (FAT) allows location of each sector of the file bystoring a chain of pointers for the file. Each pointer points to thenext sector of the file.

Hard drives suffer from a number of disadvantages. Their size, theirheight in particular, often makes them unattractive for use in portableand lap top computers. The height of a hard disk drive has often been alimiting factor in attempts to reduce computer size to make computersmore portable. Hard drives also consume relatively large amounts ofpower, which makes them even less attractive for computers that arebattery powered.

Hard drives are less than ideal for use in computers that will be usedout-of-doors. Magnetic disk drives are extremely sensitive to vibrationand shock. Additionally, magnetic drives do not tolerate well the dustand humidity associated with much of the great outdoors.

Semiconductor memories, also referred to as solid state memories, do notsuffer from many of the disadvantages of magnetic disk drives. They aresmall and rugged and consume significantly less power than magneticdrives.

One type of non-volatile semiconductor memory is the FLASH electricallyerasable programmable read only memory (FLASH memory). FLASH memoriescan be programmed by the user and once programmed retain their datauntil erased. FLASH memories are erased by erasing a block of memory ata time. A block is a relatively large amount of data, 64 Kbytes or65,535 bytes.

A FLASH memory cell is erased when the net charge on the floating gateis neutral. An erased FLASH bit is referred to as a "1." Conversely, aFLASH memory cell is programmed when the net charge on the floating gateis negative. A programmed FLASH bit is referred to as a "0." A FLASHmemory cell cannot be reprogrammed without a prior erasure with oneexception. Any FLASH bit can be programmed to a 0 at any time.

In some prior FLASH memories the programming and erasure is controlledinternally by a write state machine or a command register. Internalcontrol of programming and erasure lessens the knowledge and timerequired to program and erase the FLASH memories. However, FLASH erasetime times remain relatively slow despite internal control. Erase cycletime is on the order one to two seconds. If an erase must occur beforeevery write or re-write of a sector of data it is not possible toapproach magnetic disk drive write times using FLASH memory.

SUMMARY OF THE INVENTION

A method of locating a sector of data stored in a solid state memorydisk given a sector numbers and a pointer to a block sector translationtable is described. A header located near the pointer is selected andexamined. That header is the desired header for the sector of data to belocated if the header stores a logical sector number equal to the givensector number. The location of the sector of data within the block canbe determined using a block sector offset stored within the desiredheader. However, if the selected header is not the desired header, thena scan count is incremented. Headers are selected and examined until thesector of data associated with a given sector number is located, or amaximum number of headers have been scanned, whichever occurs first.

Other objects, futures and advantages of the present invention will beapparent from the accompanying drawings and the detailed descriptionthat follows.

BRIEF DESCRIPTION OF THE FIGURES

The present invention is illustrated by way of example and notlimitation in the figures of the accompanying drawings in whichreferences indicate similar elements and in which:

FIG. 1 is a block diagram of a personal computer including a solid statememory disk.

FIG. 2 is a block diagram of a solid state memory disk.

FIG. 3 is a block diagram of a block file structure.

FIG. 4 is a block diagram of a first embodiment of the sector headertranslation table.

FIG. 5 is a flow diagram of an algorithm for building a sector headertranslation table upon power-up,

FIG. 6 is an object diagram of the algorithms of the solid state diskcontroller.

FIG. 7 is a flow diagram for a disk read.

FIG. 8 is a flow diagram for seeking a sector of data using a blocksector translation table.

FIG. 9 is a flow diagram for writing a sector.

FIGS. 10a and 10b are a flow diagram of a first method of allocatingmemory space within the FLASH array.

FIG. 11 is a diagram block chains.

FIGS. 12a and 12b are a flow diagram of a method of allocating memoryspace in response to a write command.

FIG. 13 is a flow diagram of a mark dirty algorithm.

FIG. 14 is a block diagram of the various services within clean-up.

FIG. 15 is a flow diagram for enabling a clean-up state machine.

FIG. 16 is a flow diagram for executing one state of a clean-up statemachine.

FIG. 17 is a flow diagram of the states of a clean-up state machine.

FIG. 18 is a flow diagram of the algorithm for evaluating whetherclean-up should be triggered.

FIG. 19 is a flow diagram of the algorithm for adaptively triggeringclean-up.

FIG. 20 is a flow diagram of the algorithm for choosing a block toclean-up.

FIG. 21 is a flow diagram for allocating memory for clean-up.

FIG. 22 is a flow diagram of a method of allocating free physical memoryfor clean-up.

FIG. 23 is a flow diagram for copying a sector into the sector buffer.

FIG. 24 is a flow diagram for copying part of a sector out of the sectorbuffer and into a destination block.

FIG. 25 is a flow diagram for updating databases after copying a sector.

FIG. 26 is a flow diagram for initiating erasure of a block.

FIG. 27 is a flow diagram for determining whether erasure of a block iscomplete.

FIG. 28 is a flow diagram for updating databases after block erasure.

FIG. 29 is a flow diagram for foreground clean-up.

FIG. 30 is a flow diagram for forcing clean-up.

FIG. 31 is a flow diagram of another way of forcing clean-up.

DETAILED DESCRIPTION I. Overview of Solid State Disk

FIG. 1 illustrates in block diagram form personal computer 50. Personalcomputer 50 includes central processing unit (CPU) 52 and monitor 54 forvisually displaying information to a computer user. Keyboard 56 allowsthe computer user to input data to CPU 52. By moving mouse 58 thecomputer user moves a pointer displayed on monitor 54. Personal computer50 uses solid state memory disk 60 for mass memory storage, rather thana hard magnetic disk. Solid state disk 60 includes solid state diskcontroller 64 to control nonvolatile semiconductor memory array 62.Nonvolatile semiconductor memory array 62 is also referred to as FLASHarray 62.

Sector data associated with a sector number is not stored at a fixedphysical location within FLASH array 62. The present invention providesa method of locating a sector of data given its sector number. Brieflydescribed, sectors of data are located using a two tiered approach.First, a pointer into a block sector translation table is obtained froma sector header translation table using a sector number associated withthe sector of data. Second, headers near the pointer in the block sectortranslation table are examined to see if any header includes a logicalsector number equal to the given sector number. Headers are examineduntil the desired header is located or a maximum number of headers havebeen scanned, whichever occurs first. Given the desired header, thesector of data is located using a block sector offset stored within thedesired header.

A. FLASH Array and Block File Structure

A greater understanding of solid state disk controller 64 is aided by anunderstanding of the object of its control, FLASH array 62. The programand erase characteristics of FLASH array 62 strongly influence solidstate disk controller 64. The FLASH devices within FLASH array 62 mustbe erased a block at a time, but can be programmed a byte at a time.Once programmed to a 0, a bit of FLASH memory cannot be programmed to a1 without first erasing an entire block. Erased bytes of memory arereferred to as "free" because they are ready to be written.

Erasure of FLASH memory is a slow process. Performing an erase each timea sector of data is written is impractical. Writes would be slow andpower consumption inefficient because an entire block, 128 Kbytes, wouldhave to be erased just to write one sector, 512 bytes. To allow rapidsector writes, solid state disk controller 64 writes each sector of datato a new, free location each time a sector is written. A result of thiswrite method is that there may be several versions of the sector dataassociated with a single sector number. The most recent version of thesector data is referred to as a "good sector," "valid sector" or a "usersector." In contrast, the earlier version of the sector is invalid andwill be marked as "dirty."

The actual amount of FLASH memory within FLASH array 62 cannot greatlyexceed the amount stated as available to the user because FLASH memoryis relatively expensive. Stated another way, when the amount of reservememory within FLASH array 62 is lean dirty sectors must be convertedinto free memory to ensure the availability of memory for writes.

FIG. 2 illustrates in block diagram form FLASH array 62 and solid statedisk controller 64. In one embodiment, FLASH array 62 uses thirty 1megabyte, by 8, FLASH memory chips. These FLASH memories include a writestate machine for automatically controlling erasure and programming.These thirty FLASH memory devices function as a 40 megabyte memory diskwhen data compression is used. Each FLASH chip inputs and outputs data 8bits at a time. To permit word-wide input and output, FLASH array 62 isorganized as pairs of FLASH devices, only one chip pair 66 of which isshown. High chip 68 of chip pair 66 stores the high byte of a word,while low chip 70 stores the lower byte of a word. Solid state diskcontroller 64 is thus able to treat each chip pair as a single 16bit-wide memory device. Word-wide input and output gives solid statedrive 60 a speed advantage compared to magnetic drives, which use serialbit stream I/O.

Each chip pair is organized as 16 blocks, each including 128 Kbytes ofmemory. Because each block of memory can store many sectors of data,each block includes, as illustrated in FIG. 3, a block sectortranslation table (BSTT) 84 to identify and locate each sector of data.

FIG. 3 illustrates block 80 and the file structure used by it and allother blocks. Block 80 is represented as a single word wide structurebut is actually stored in two FLASH chips. The high byte of each word isstored in high chip 68 and the low byte of each word is stored in lowchip 70.

The data structure of block 80 includes block sector translation table84 and data space 86. Block sector translation table 84 stores headers.A header is a block of information about one logical sector number andits associated data. As used herein a logical sector number (LSN) refersto a sector number stored within a BSTT. A sector number is a sectoridentifier received from CPU 52, which the CPU believes corresponds to afixed physical location. However, as a result of the write policy usedby solid state disk 60, an LSN does not correspond to a fixed physicallocation. Also as a result of the write policy used, several headers andLSNs may correspond to a single sector number.

In one embodiment, each logical sector number is 24 bits long.

A header is created for each and every sector number during diskformatting. This allows the loss of sectors of data to be detectedduring the execution of read and write commands. Failure to find theheader associated with a particular sector number indicates that theassociated sector of data has been lost. As used herein "lost" refers toa sector of data that disappears because of a defect in FLASH 62 arrayor to a sector of data that is unreliable because it has been corrupted.

Each header 85 includes a cyclical redundancy check (CRC), which allowssolid state disk 60 to determine the reliability of header 85.

Header 85 also includes an attribute word that contains a great deal ofinformation about the sector data associated with the header. One bit ofthe attribute word indicates whether the sector number has been markedas part of a bad track. Another bit indicates whether or not the sectordata has been compressed. The attribute word includes two dirty bits forreliability. The sector data associated with the header is consideredvalid if both dirty bits are set and dirty if either dirty bit is reset.The attribute word also includes a data attached bit. When no sectordata is attached to the header, the data attached bit is reset. This isthe case for all headers after formatting. Once data is written for thesector number, the data attached bit is set. The final piece ofinformation included in the attribute word is a revision number. Therevision number allows solid state controller 64 to identify the validheader when multiple valid headers with the same LSN exist.

The last piece of information stored in header 85 is a block sectoroffset (BSO). The BSO is an offset from the top of the block to thestart of FLASH memory space associated with the header. Memory space isallocated to a header whether or not data is stored in that space. Ifdata is not attached to the header, then the amount of memory allocatedis the size of data space 86 divided by the maximum number of headers inBSTT 84.

Memory space must be allocated to headers that do not have sector dataattached. This is because during formatting entire BSTTs 84 are filledwith headers that have no data attached. Even though all of data space86 associated with these BSTTs 84 is free, data space 86 cannot bewritten to. There is simply no room in the BSTTs of these blocks. Toavoid a Catch 22 situation in which free memory is not available forallocation, all of data space 86 associated with these BSTTs 84 iscategorized as user data. In other words, the sector is neither dirty,as indicated by dirty bits, nor free. By allocating memory data space 86to headers with unattached data and by subtracting that amount of memoryfrom the amount of free memory in the appropriate block, appropriatechip, and array, part of data space 86 is designated as user data. Assectors of data are written after formatting, the headers withoutattached data will be marked dirty, along with their allocated dataspace. Eventually, blocks filled up during formatting will be cleaned-upand their data space made available as free memory.

It is crucial that headers can be written without data attached whendata compression is not used. Few reserves are available for clean-upwhen data compression is not used and every possible sector of data iswritten to solid state disk 60. In this situation, clean-up efficiencyis seriously impaired, and data fragmentation is likely. Datafragmentation, in turn, leads to many foreground clean-ups, degradingoverall solid state disk performance.

Because data for one sector is butted up against data for another sectorin data space 86, each BSO indicates the top of data for one sector andthe bottom of data for another sector. For example, the block offset forsector 1, BSO₁, points to the start of the data associated with LSN₁.BSO₁ also points to the bottom of data associated with LSN₂. Another wayof looking at block sector translation table 84 is that each LSN issandwiched between two block sector offsets, one pointing to the top ofthe sector data and the other pointing to the bottom of the sector data.The bottom of the data associated with the sector stored at the bottomof the block, LSN₁, need not be indicated by a BSO as the location ofthe bottom of the block is known.

Block sector translation table 84 grows down toward data space 86. Thefirst header BSTT 84 is written at the top of block 80. The header foreach succeeding sector written into block 80 is stored below theprevious header.

In contrast to BSTT 84, data space 86 grows upward. The first sector ofdata written into block 80 is written into the bottom of data space 86.The next sector of data written into data space 86 is writtenimmediately above the previous sector. For example, the data associatedwith LSN₂ is located within a lower range of addresses than the dataassociated with LSN₁.

The number of sectors that may be stored in data space 86 before it runsinto BSTT 84 varies. This variability arises when sector data iscompressed, which causes sector size to vary. With data compression,sector size may vary between 256 words to just a few words. A maximum of2047 sectors may be stored in data space 86 given the seek strategy usedby solid state disk 60.

Solid state disk controller 64 prevents BSTT 84 and data space 86 fromcrashing into each other as they grow. In fact, solid state diskcontroller 64 ensures that some slack 88 is maintained between BSTT 84and data space 86. Slack 88 is free FLASH memory, which has not beenprogrammed. According to the conventions of FLASH memory, a free memorylocation stores FFFF (hexadecimal). During seeks of block sectortranslation table 84, slack 88 indicates that the end of BSTT 84 hasbeen reached.

Block 80 also stores block attribute data 90. Information specific tothe block is stored attribute data 90. For example, block attribute data90 includes cycle count, which indicates the number of times the blockhas been erased and written to. Block attribute data 90 may also includeblock defect information.

B. Overview of the Solid State Controller Hardware

Referring once again to FIG. 2, reading, writing, and clean-up of FLASHarray 62 is controlled by solid state disk controller 64. Microprocessor92 manages these tasks using database 93, sector header translationtable (SHTT) 94 and the algorithms of the present invention, which arestored in FLASH memory 95. Databases 93, SHTT 94, and algorithms will bedescribed in greater detail below. An application specific integratedcircuit, window ASIC 96, serves as a window to the outside world formicroprocessor 92 as well as windows into the FLASH array 62. Via windowASIC 96, microprocessor 92 receives commands from CPU 52 and reads andwrites to FLASH array 62. Window ASIC 96 includes four windows, allowingmicroprocessor 92 to repeatedly and rapidly access a number of FLASHmemory locations. Data to and from CPU 52 flows through window ASIC 96to sector buffer 98. Sector buffer 98 allows data transfers to CPU 52 tooccur more rapidly that possible otherwise. Solid state disk controller64 includes charge pump 100. Charge pump 100 converts 5 volts into the12-volt level needed for programming and erasure. Charge pump 100 is notnecessary unless the voltage levels supplied to solid state disk 60 areall below the 12-volt voltage level necessary to program the FLASHdevices within FLASH array 62.

C. FLASH Array Database and Sector Header Translation Table

Performance of solid state disk 60 is enhanced by storing repeatedlyused information in random access memory (RAM). This information isreferred to as FLASH array database 93 because it generally relates tocurrent characteristics of FLASH array 62. Information within FLASHarray database 93 includes:

1. The total number of dirty words within FLASH array 62 (TDFA);

2. The total number of dirty words in each block (TDFB_(N));

3. The total number of free words within FLASH array 62 (TFFA);

4. The number of free words within each chip pair (TFFC_(M)); and

5. The total number of free words within each block (TFFB_(N)).

The use of FLASH array database 93 by solid state disk controller 64will be described below on an algorithm by algorithm basis.

Sector header translation table (SHTT) 94 translates a sector numberinto a pointer to an associated sector of data. To permit its frequentand easy modification, SHTT 94 is preferably stored in RAM. The pointersupplied by sector header translation table 94 does not point directlyto the data associated with the given sector number. Rather, the pointersupplied by SHTT 94 points to a location within BSTT 84 near the headerassociated with the sector number.

The reason why SHTT 94 does not point directly to the header associatedwith the sector number can be understood by examining the number of bitsnecessary to uniquely address each possible header in a BSTT 84. Amaximum of 2047 headers may be stored in a BSTT 84. Given 15 chip pairs,each including 16 blocks in a chip pair, FLASH array 62 could store asmany as 491,280 headers. Uniquely identifying that many headers requires19 bits. Storing 19 bits requires using three by 8 RAM chips. Thus, 5 of8 bits in one RAM would be wasted to store three bits. This is anexpensive and unacceptable solution in the effort to produce a pricecompetitive solid state memory disk.

A number of SHTT designs solve this 19 bit dilemma. One embodiment,shown in FIG. 4, stores a pointer for every one of the possible 83,300sectors numbers of data of an industry standard 40 MB disk drive. Only16 bits are stored for each entry in SHTT 94. Four of the bits indicatethe chip pair in which the sector data being sought is stored. Anotherfour bits indicate the particular block in which the sector is stored.The remaining 8 bits represent an offset from the top of the block toheader which is near the desired header. This offset is called a headerpointer. Using 8 bits to represent the header pointer means that thereare 256 header pointers available to locate a maximum of 2047 headers.Consequently, a maximum of 8 headers must be scanned in BSTT 84 tolocate the desired header. Stated another way, SHTT 94 requires a twotiered search to locate the data associated with a particular sectornumber. First, SHTT 94 is searched to locate a particular chip, blockand header pointer. Second, the indicated area of BSTT 84 is searched tofind the header associated with the sector number.

SHTT 94 of FIG. 4 stores a physical address for each sector stored onsolid state disk 60. The number of entries in SHTT 94 can be reduced byusing a most recently used cache memory.

Both FLASH array database 93 and SHTT 94 must be generated duringpower-up because they are stored in volatile memory, RAM, and becausereads and writes depend upon these 93 and 94. FIG. 5 illustrates analgorithm to build both FLASH array database 93 and SHTT 94. Using thisalgorithm both SHTT 94 and the FLASH array database 93 are generated byscanning each BSTT 84. The location of the header associated with eachsector number is noted, as well as the amount of free and dirty memorywithin the block.

Building begins in step 110 with the initialization of SHTT 94. Eachpointer for each sector number is set to the same initial value. In oneembodiment, that value is FFFF (hexadecimal). As a result, afterbuilding SHTT 94 retrieving a pointer equal to the initial valueindicates that the sector of data associated with the sector number hasbeen lost because during formatting a header is created for each sectornumber. Microprocessor 92 branches from step 110 to step 111.

In step 111 total amount of free memory within FLASH array, TFFA, andthe total amount of free memory per chip, TFFC_(M), are initialized totheir maximum values. The maximum free FLASH per chip is 64Kbyte/block * 16 blocks, or 2048 Kbyte per chip pair minus the blockattribute data stored in each block. Similarly, TFFA_(max) is 2048 Kbyteper chip pair * 15 chip pairs, or approximately 30 Mbyte. This done,microprocessor 92 branches to step 112 to begin scanning BSTTs 84.

In step 112 microprocessor 92 determines whether any block within FLASHarray 62 remains to be scanned. If so, microprocessor 92 proceeds tostep 114 and selects a block to scan, B_(C). From step 114,microprocessor 92 branches to step 115. There the total amount of freememory in the current block, TFFB_(C), is initialized to its maximum.Microprocessor 92 then proceeds to step 116.

In step 116, microprocessor 92 reads the next header from the selectedblock's block sector translation table 84. Afterward, microprocessor 92branches to step 118.

Microprocessor 92 begins to classify the current header in step 118.Microprocessor 92 determines whether the header has been marked dirty byreading the dirty bits included in the attribute word of the currentheader. If either dirty bit is a logical zero, the data associated withthe LSN is considered dirty. If the sector of data is dirtymicroprocessor 92 branches to step 120 to update FLASH array database93.

In step 120, microprocessor 92 determines the size of the dirty sector,which may vary as a result of data compression. Sector size isdetermined by comparing the block sector offsets on either side of thecurrent LSN. After determining the size of the dirty sector,microprocessor 92 updates TDFA and TDFB_(C). The total number of dirtywords, TDFA, and the total number of dirty words within the currentblock, TDFB_(C), are increased by the size of the sector and its header.Afterward, microprocessor 92 advances to step 121.

Microprocessor 92 appropriately revises its FLASH free variables. Thesize of the current sector and its header are subtracted from TFFA,TFFB_(C) and TFFC_(C). Afterward, microprocessor returns to the buildingof FLASH array database 93 and SHTT 94 by branching back to step 116.

On the other hand, if the sector is not dirty microprocessor 92 branchesfrom step 118 to step 122. Microprocessor 92 then examines the attributeword within the current header to determine if the end of the BSTT 84has been reached. If the attribute word for the next header is FFFF(hexadecimal), the end has been reached. In this case, microprocessor 92branches back to step 112 to determine whether scanning of BSTTs 84should continue. In all other cases, microprocessor 92 advances to step124 from step 122.

Entry into step 124 means that the current header is associated with avalid sector of user data. Accordingly, microprocessor 92 reduces freevariables TFFA, TFFB_(C) and TFFC_(C) by the size of the current sector.(The subscript "C" designates "current.") This done, microprocessor 92advances to step 126.

In step 126, revision of the SHTT entry for the current header begins byseeking an entry in SHTT 94 for the sector number equal to the LSNincluded in the current header. If the current entry in the SHTT is FFFF(hexadecimal), no information exists for the sector number.Microprocessor 92 responds to this situation by branching to step 128.Any entry other than FFFF for the sector number means that there are twovalid versions of the data for that sector number. To sort things out,microprocessor 92 advances to step 130.

Consider first the simpler situation; that is, when no data is stored inSHTT 94 for the sector number corresponding to the current LSN. In step128, microprocessor 92 writes the chip, block and header pointer for thecurrent header into SHTT 94. This done, microprocessor 92 branches backto step 116 to examine another header.

Things are more complicated when SHTT 94 already includes an entry forthe sector number associated with the current LSN. In step 130, themicroprocessor determines which header and thus which sector of data ismost current by comparing their revision numbers. In step 132,microprocessor 92 then marks dirty the earlier, invalid, sector of databy programming the dirty bits in its header to a logical 0.

In step 133, the microprocessor updates database 93 to reflect the totaldirty for the block including the sector just marked dirty and for FLASHarray 62. The update is accomplished by adding the sector size toTDFB_(C) and TDFA.

SHTT 94 is finally updated, if necessary, in step 134. No update isnecessary if the header and LSN selected in step 116 were marked dirtyin step 132. Otherwise, the chip, block and header pointer for the LSNselected in step 116 are written into SHTT 94.

From step 134, microprocessor 92 returns to step 116. There anotherheader is selected.

After every header within FLASH array 62 has been scanned, building ofSHTT 94 is complete. If a header has not been located which correspondsto a sector number, the pointer within SHTT 94 remains at its initialvalue. Thus, the pointer indicates that the sector of data has been lostand every sector of data is accounted for by SHTT 94.

II. Solid State Disk Controller

The heart of solid state disk controller 64 is the set of algorithmsstored within FLASH memory 95. These algorithms control the reading,writing, and cleaning-up of FLASH array 62. These algorithms help createthe illusion that CPU 52 is dealing with an industry-standard hardmagnetic disk drive.

The object diagram of FIG. 6 illustrates the general organization andinterrelationship of algorithms used by solid state controller 64. Thealgorithms of solid state disk controller 64 are organized into threeclasses: top level scheduler 150, host interface 152, and FLASH media154. Top level scheduler 150 handles the allocation of microprocessor 92processing time between the other two classes 152 and 154. Hostinterface 152 interprets industry standard disk drive commands from CPU52 and translates them into commands that FLASH media 154 can act upon.FLASH media 154 interfaces directly with FLASH array 62, responding toread and write requests from host interface 152. FLASH media 154 alsomanages the cleaning-up of FLASH array 62.

The scheduling of host interface 152 and FLASH media 154 is a relativelysimple task. Power-up of solid state disk 60 fires top level scheduler150. It initializes the solid state disk 60 and then calls hostinterface 152. This allocates to CPU 52 all the resources of solid statedisk 60. When host interface 152 returns control to top level scheduler150, clean-up object 164 is called. If a clean-up is on-going, then aslice of CPU execution time, 500μ seconds in one embodiment, isallocated for clean-up. When clean-up returns control to top levelscheduler 150, host interface 152 is called again. Top level scheduler150 repeats the process again and again for as long as solid state disk60 is powered up.

Host interface 152 includes two classes of algorithms, interface 156 andsector buffer 158. Interface 156 emulates an AT-IDE hard disk interface.Interface 156 handles ATA command interrupts and translates ATA commandsinto commands comprehensible by FLASH media 154. In alternateembodiments, host interface 152 may simulate a SCSI disk interface orother standard disk drive interface. Sector buffer 158 manages the usageof sector buffer 98.

Briefly described, FLASH media 154 includes five types of algorithms, orservices: disk 160, sectors 162, clean-up 164, FLASH power 166 and FLASHarray interface 167. Disk 160 services read and write requests frominterface 156. Disk 160 also translates other commands from interface156 and delegates them for execution to fellow classes 162, 164, 166 and167. Sectors 162 is responsible for most tasks relating to sectorsstored within FLASH array 62. Sectors 162 maintains FLASH array database93 used by both disk 160 and clean-up 164, as well as allocating freememory space within FLASH array 62. Sectors 162 also controls thereading of data from FLASH array 62 via FLASH array interface 167 andsequencer 168. FLASH power 166 is essentially a device driver; that is,it generates the voltage levels necessary to read or write to FLASHarray 62. FLASH array interface 167 handles the low level routines whichdirectly control the reading and writing of FLASH array 62. Sequencer168 handles the movement of data between sector buffer 98 and FLASHarray 62. Just as its name implies, clean-up 164 manages the clean-up ofFLASH array 62. FLASH power 166 manages the limited current budget ofsolid state disk 60, which in one embodiment is only 120 mA. Given thaterasing a block in a chip pair requires up to 60 mA efficient managementof current is a concern.

A. Reading a Sector from FLASH Array 62

Briefly described, reading a sector is a three step process. First, SHTT94 is searched for a pointer to the header associated with the sectornumber. Second, the header is located and its attribute word is examinedto see if the attached data is valid. Third, if the sector dataassociated with header is valid, its location is determined and returnedto CPU 52.

A read algorithm implemented by disk 160 is illustrated in the flowdiagram of FIG. 7. All callers input a sector number.

The first task of microprocessor 92 in step 170 is examination of theinput sector number for validity. The maximum number of sectors withinan industry standard disk drive determines whether a sector number isvalid. For example, a standard 40 MB drive includes 83,300 sectors.Thus, any sector number greater than 83,300 would be invalid in a 40 MBdrive.

Microprocessor 92 branches from step 170 to step 172 if the callingprogram passed in an invalid sector number. Microprocessor 92 indicatesthe invalidity of the sector number to the caller by setting a statusword to invalid sector number range. Microprocessor 92 then branches tostep 174, returning control to the caller.

On the other hand, microprocessor 92 proceeds to step 176 from step 170if the caller has requested to read a sector with a valid sector number.In step 176, microprocessor 92 takes the sector number and searches SHTT94 for the chip, block and header pointer associated with the sectornumber.

Microprocessor 92 determines in step 178 whether data exists for thesector number requested. Microprocessor 92 may be able to tell whether aheader exists for a particular sector number by examining the headerpointer retrieved from SHTT 94. If the header pointer is FFFF(hexadecimal), a valid header cannot be found for the sector number,even though a header is created for every sector number duringformatting.

Microprocessor 92 responds to a lost sector by branching to step 180from step 178. There microprocessor 92 indicates that the header was notfound. Microprocessor 92 then branches to step 174, the read complete.

On the other hand, if a header exists for the sector number thenmicroprocessor 92 branches to step 182. There sectors 162 seeks for theheader associated with the input sector number. If the desired header islocated, sectors 162 indicates the chip, block, and offset to thatheader. Sectors 162 also indicates if the desired header cannot befound.

Once sectors 162 returns from its seek, microprocessor 92 branches tostep 184 to determine the success of the seek. If the seek algorithm didnot locate the header associated with the sector number, microprocessor92 branches to step 180 to so inform the caller. On the other hand, ifthe appropriate header was located microprocessor 92 branches to step186. There the information retrieved in step 182 is passed to the readservice of sectors object 162. The data associated with the sectornumber will be copied into sector buffer 98 and provided to host CPU 52.Afterward, the read complete, microprocessor 92 returns control to thecaller by branching to step 174.

FIG. 8 illustrates a method of seeking sector data given a sectornumber. This algorithm also is implemented by sectors 162.

Briefly described, the seek algorithm of FIG. 8 has two approaches tolocate the data associated a sector number. The first approach reliesupon data coherency to decrease seek time. That is, it is assumed thatthe sector of data being currently sought is related to, and locatednear, the last sector of data located. If this approach fails to locatethe desired sector of data, then a two tiered approach is used. First,SHTT 94 is searched to located a chip, block and a header offset for thesector number input, SN_(i). Then, using that information, a BSTT 84 issearched for a header with an LSN equal to the sector number input bythe caller.

Location of the sector of data begins in step 190. There microprocessor92 determines whether another header is stored in the last BSTT 84searched. If so, it is possible that the header associated with inputsector number resides within the same BSTT 84. In that case,microprocessor 92 advances to step 192. However, if the last headerlocated was the last header located within its BSTT 84, then it is notpossible to quickly locate the sector of data associated with the inputsector number. A two tiered search must be used. Microprocessor 92advances to step 206 in that case to begin that two tiered search.

Consider first, the actions of microprocessor 92 upon entry to step 192.Microprocessor 92 reads the header immediately following the last headerlocated. Let us designate the LSN in this header "LSN_(L+1)." LSN_(L+1)is compared to SN_(i) in step 194 to determine if the current header isthe desired header. If LSN_(L+1) does not equal SN_(i) then the sectorof data must be located another way. In that case, microprocessor 92advances to step 206. On the other hand if LSN_(L+1) does equal SN_(i)then microprocessor 92 branches to step 196.

With step 196 the process of determining the reliability of the currentheader begins. During step 196 microprocessor 92 generates a cyclicalredundancy check (CRC) for the current header. Let us call this CRC"CRC_(G)." In step 198 CRC_(G) is compared to the CRC stored in theheader itself, CRC_(stored). The information within the current headeris considered reliable when CRC_(G) =CRC_(stored). Microprocessor 92responds to this condition by branching to step 200. On the other hand,if the current header is not reliable because CRC_(G) does not equalCRC_(stored) then microprocessor 92 proceeds to step 206.

Consider first the actions of microprocessor 92 when the desired headerhas been located. Microprocessor 92 first determines the size of thesector of data associated with the header in step 200. This is doneusing the appropriate block sector offsets, as previously discussed. Thesector size, chip, block and offset to the desired header are returnedto the caller in step 202.

When it is successful, the method of steps 190-202 reduces the timerequired to locate a sector of data by a factor of four as compared tothe two-tiered search of steps 206-222.

The two-tiered search begins with step 206 after the quicker approachhas failed. Microprocessor 92 performs the first level of search byseeking SHTT 94 for the chip, block and header pointer associated withthe sector number input SN_(i). The values for the chip, block andheader pointer retrieved in step 206 are examined in step 208 todetermine whether a header can possibly be located for SN_(i). If thevalues are equal to the initial value, then the header associated withSN_(i) has been lost. In one embodiment the initial, invalid values areFFFF (hexadecimal). Microprocessor 92 responds to invalid chip, blockand header pointer values by branching to step 224. There the caller isinformed that the header associated with input sector number was notfound. On the other hand, if the values retrieved from SHTT 94 arevalid, microprocessor 92 advances to step 210.

The second tier of the seek begins in step 210 by initializing a scancount to zero. The scan count tracks the number of header examinedduring the seek. Microprocessor 92 then reads header near the headerpointer.

From step 210, microprocessor 92 advances to step 212 to compare thecurrent scan count to the maximum scan count.

The maximum number of headers that must be scanned is set by dividingthe maximum number of headers in a BSTT 84 by the maximum number ofheader pointers. For example, in one embodiment of solid state disk 60,a maximum of 2047 headers are stored in BSTT 84 and only 256 headerpointers are used. Thus, a maximum of 8 headers must be scanned in thisembodiment. If the scan count is less than the maximum, microprocessor92 branches to step 214 to continue the search for the desired header.On the other hand, once the scan count equals the maximum microprocessor92 advances to step 224 from 212.

Consider first the situation when the scan count is less than themaximum. Microprocessor 92 enters step 214 where it compares the LSNstored in the current header to the input sector number. If the two areequal the desired header may have been located. Otherwise, the searchfor the desired sector must continue.

The search of the desired sector continues in step 222 by incrementingthe scan counter and reading another header located beneath the lastheader. Microprocessor 92 then returns to step 212 to determine if themaximum scan count has been reached and then step 214 to determinewhether this new header is the desired header.

When an LSN is equal to the input sector number, microprocessor 92advances to step 216 from step 214. In step 216 a cyclical redundancycheck, CRC_(G), is generated for the current header. CRC_(G) is used instep 218 to assess the reliability of the information included in theheader.

There microprocessor 92 determines whether the current header is the onesought by comparing the sector number input, SN_(i), to the LSN storedin the block. If LSN=SN_(i), microprocessor 92 advances to step 216.

In step 218 reliability of the header is determined by comparing CRC_(G)to the CRC stored in the selected header. The information within theheader cannot be relied upon unless the two CRCs are equal.

Microprocessor 92 branches to step 220 from step 218 if the currentheader is not reliable. Because the block sector offset within thecurrent header cannot be relied upon, two sectors of data are lost. Forexample, assume that the header being scanned was the header for LSN₁,illustrated in FIG. 3. Because BSO₁ cannot be relied upon, neither thestart of data for LSN₁, nor the end of data for LSN₂ can be determined.Thus, in step 220 microprocessor 92 discards two sectors of data bymarking dirty both the selected header and the header immediatelybeneath it. Microprocessor 92 then advances to step 222.

If the current header can be relied upon, microprocessor 92 advances tostep 200 from 218. Having reached step 200, microprocessor 92 beginsgathering the information needed to read the sector data. First, thesize of the sector is determined by reading the appropriate BSOs.Finally, in step 202 the sector size and attribute word is returned tothe caller along with the chip, block and offset to the header.Microprocessor 92 then returns control to the caller in step 204.

B. Writing a Sector to FLASH Array 62

Writing a sector of data involves three major tasks. First, enough freememory to write the sector data must be located and reserved. Second,any previous version of sector data with the same LSN is marked dirtyand sector header translation 94 is updated. Third, the sector data andits header are written into the appropriate block.

FIG. 9 presents an overview of the process of writing a sector of datainto FLASH array 62. A write begins when disk 160 receives a writecommand accompanied by an input sector number, SN_(i), from CPU 52. Atthis point, the data to be written into FLASH array 62 resides withinsector buffer 98. A number of outcomes are possible using the algorithmof FIG. 3 depending upon the LSN_(i) and the data stored in SHTT 94.

Consider first the outcome when SN_(i) is invalid. The first action ofmicroprocessor 92 upon receipt of the write command is to examineSN_(i). This is done in step 230. Any SN_(i) greater than 83,300 or lessthan 1 will not be written into a 40 MB drive.

Given that the SN_(i) is invalid, microprocessor 92 branches from step230 to step 232. There microprocessor 92 sets the status word to informinterface 156 that the sector number received was invalid. That done,microprocessor 92 advances to step 234. Control of microprocessor 92then returns to the calling program.

Assume now that SN_(i) is valid, and that its attribute word indicatesthat the track associated with SN_(i) has been marked bad. BecauseSN_(i) is valid, microprocessor 92 branches to step 236 from step 230.

Preparations to write the sector begin in earnest in step 236.Microprocessor 92 enables charge pump 100, enabling it to develop the 12volt level necessary to program FLASH memory. Microprocessor 92 thenperforms other tasks until charge pump 100 reaches 12 volts.

In step 238, microprocessor 92 calls the seek algorithm to locate thechip, block and header pointer for an earlier version of the sectordata. After locating the information associated with the sector number,microprocessor 92 branches to step 240.

In step 240, microprocessor 92 determines whether the bad track bit hasbeen set. Microprocessor 92 does so by examining the sector attributeword. Given that the bad track bit has been asserted, microprocessor 92advances to step 242 from step 240.

Microprocessor 92 will not write the sector data into FLASH array 62because the track associated with SN_(i) has been marked bad. In step242, microprocessor 92 informs interface 156 that the track is bad.Afterward, microprocessor 92 returns control to interface 156 bybranching down to step 234.

The previous discussions show that a sector of data will not be writteninto FLASH array 62 if either SN_(i) is invalid or its associated trackhas been marked bad. Stated conversely, sector data will be written intoFLASH array 62 if it is associated with a good track and valid SN_(i).Consider now such a situation. Microprocessor 92 responds by branchingfrom step 240 to step 244.

In step 244, microprocessor 92 continually queries charge pump 100 todetermine whether it has reached 12 volts. The amount of time requiredby charge pump 100 to reach 12 volts cannot be predicted because itvaries depending upon how long it has been since charge pump 100 waslast enabled. The less time that has passed since charge pump 100 wasfired, the less time required to reach 12 volts again. Once charge pump100 stabilizes at 12 volts, microprocessor 92 branches to step 246.

In step 246, microprocessor 92 determines whether a previous header withthe same LSN should be marked dirty. Microprocessor 92 makes thisdetermination based upon the information retrieved by the seek of step238. If a header was located, microprocessor 92 proceeds to step 247 tomark that header dirty. Afterward, microprocessor 92 advances to step248 to determine whether the previous header was marked dirty or whetherthe task was cached. If task was not cached, microprocessor 92 advancesto 250. Otherwise, microprocessor 92 branches to step 249. Because themark dirty was cached, there will be two headers with the same LSN atthe end of the current write. To distinguish the valid data afterpower-loss, the revision number for the LSN associated with the mostcurrent version is incremented in step 249. Microprocessor 92 thenproceeds to step 250.

With step 250, microprocessor 92 begins the process of writing the newversion of the sector data within FLASH array 62. Microprocessor 92allocates sufficient free memory within FLASH array 62 to store thesector of data and header. This is an involved process that will bedescribed in detail later. Suffice it to say that allocation of memoryrequires locating sufficient memory within data space 86 of a block andmarking that memory space as reserved. Microprocessor 92 then exits tostep 252.

Microprocessor 92 completes the writing of the header in steps 252 and254. First, in step 252, a CRC is generated for the header, whichexcludes the dirty bits and revision numbers because they may be changedin the course of events. Afterward, in step 254, the CRC, attribute wordand LSN are written into BSTT 84. The LSN is set equal to SN_(i).

Microprocessor 92 finally writes the sector data into data spacereserved in step 256. An error correction code, ECC, is also writtenwith the data.

The new version of the sector data safely written, in step 258microprocessor 92 updates sector header translation table 94 so that itpoints to the most recent version of the sector data associated with thesector number.

In step 260 microprocessor 92 disables charge pump 100. Turning chargepump 100 off when it is not needed reduces power consumption. Thisenhances the suitability of solid state disk 60 for use in portable andlaptop computers. This done, microprocessor 92 branches to step 262.

During step 262, microprocessor 92 forces an evaluation of whetherclean-up is necessary as a result of the write. Microprocessor 92 doesthis by calling a service named "Enable CSM," which will be described indetail later. This done, microprocessor 92 advances to step 234,returning control to the caller.

C. Allocating Memory Space for a Write

Allocating memory space within FLASH array 62 is a complex and criticaltask. Not only must memory space be allocated, an appropriate locationmust be chosen to prevent performance degradation. Choosing a block fora sector write potentially involves four major decisions. First, arethere sufficient FLASH memory reserves to allow the write? Second, isthere enough free memory in the block to which the current process waslast allocated to store the current sector? If the answer to the secondquestion is no, then a third and a fourth question must be asked. Isthere a block with enough free FLASH memory to store the sector data? Isthat block an appropriate block in which to store this sector?

The fourth decision is a difficult one because whether a block is anappropriate block depends on a number of factors. First, the blockchosen cannot be included in a busy chip pair. Waiting on a busy chippair is avoided because it diminishes the speed with which commands fromCPU 52 are obeyed. Second, the block chosen should not be within a blocktargeted for clean-up. Data written into a block targeted for clean-upwill just have to be relocated right away. Third, the block chosenshould not be within the same chip pair allocated to another process.This avoids data fragmentation, which eventually results in foregrounderase and decrease in power efficiency. Data fragmentation refers to arandom distribution of clean, dirty, and free sectors throughout FLASHarray 62. Data fragmentation is catastrophic to solid state diskperformance when reserves of memory space are lean. By allocating thosetargetted for CPU 52 into different chip pairs than writes initiated byclean-up, dirty sectors are likely to be grouped together due toaccumulation in those chips which are targetted for cleanup. Thisreduces the number of sectors that must be copied out of a block duringclean-up, thereby improving power efficiency and clean-up performance.

FIG. 10 is a flow diagram of an algorithm for allocating memory that maybe used for host writes or for clean-up. The algorithm shown belongs tosectors 162. Given the size of the sector to be written and the callertype, the allocate algorithm locates and reserves sufficient memory. Thereserved location is stored in RAM for future reference by thealgorithm.

A number of outcomes are possible using the allocation algorithm of FIG.10. Consider first the outcome of greatest interest: memory isallocated. Other outcomes will be considered afterward.

Allocation of memory begins in step 270 by determining whether memoryhas been requested for clean-up or for a host write. If clean-uprequires free physical memory microprocessor 92 branches ahead to step282. On the other hand, when host CPU 52 requests memory microprocessor92 must decide whether to allow the write in view of available memoryreserves. This determination takes place in steps 272 through 280.

Microprocessor 92 begins its determination in step 272 by calculatingwhat the available reserves after writing a sector of the desired sizewould be. As used here, "reserves" refer to both free and dirty memorywithin FLASH array 62 and "sector size" includes both the size of thedata and the header. The available reserves are then compared to a firstwarning level, FLASH Warn 1. In one embodiment, FLASH Warn represents 19blocks of reserves. If the available reserves exceed FLASH Warn 1,microprocessor 92 advances to step 273. There, microprocessor 92indicates that reserve status is acceptable. Afterward, microprocessor92 proceeds to step 282. On the other hand, if available reserves willnot exceed FLASH Warn 1, microprocessor 92 proceeds to examine thereserves in greater detail. First, however, microprocessor 92 signalsthat reserves have fallen below the first warning level by setting thereserve status to FLASH Warn 1 in step 274.

In step 275 the available reserves remaining after the write arecompared to a second warning level, FLASH Warn 2. In one embodiment,FLASH Warn 2 represents 11 blocks of reserve. Microprocessor 92 branchesahead to step 282 if available reserves after the write exceed thissecond warning level. Otherwise, microprocessor 92 proceeds to step 276.There the reserve status is reset to indicate reserves have fallen belowthe second warning level.

Microprocessor 92 reevaluates the available reserves in step 278 bycomparing them to a third warning level, standby. In one embodiment,standby represents 2 blocks of reserves. If available reserves willexceed standby after the write, then microprocessor 92 permits the writeby branching to step 282. Microprocessor 92 advances to step 279 ifavailable reserves will not exceed standby. There the reserve status isset to standby. As a result, no more writes by CPU 52 will be permitteduntil sectors are released by host CPU 52. Control is then returned tothe caller in step 304.

With step 282 microprocessor 92 begins searching for a block with enoughfree memory in data space 86 to store the sector data to be written.Microprocessor 92 searches for a block using a set of block chains, twoof which are illustrated in FIG. 11. Block chain 320, figurativelyspeaking, chains together block 1 of each chip pair (CP_(M)) withinsolid state disk 60. Similarly, block chain 322 links together block 15of each chip pair (CP_(M)). Blocks 2 through 14 are linked together insimilar chains, though they are not shown. Chaining blocks together inthe fashion shown is not necessary but causes sectors of data to bedistributed across chip pair boundaries. This provides passivewearleveling and nearly eliminates the need for active wearleveling.

Microprocessor 92 uses block chains in conjunction with pointers storedin RAM. Microprocessor 92 stores the chip and block last allocated toeach caller. The next time the same process attempts to write a sector,the allocation algorithm first examines the last block allocated to thatprocess. This helps keep related data "files" in contiguous memory spaceand helps reduce the possibility of data fragmentation. Asmicroprocessor 92 searches for a block for each caller, the caller'spointer circles around a block chain. The starting pointer stored in RAMfor each caller is different to help avoid data fragmentation. Eachpointer moves around the block chains at its own rate. Additionally,each pointer moves from block chain to block chain at a different rate,depending upon the availability of free memory. No two pointers everpoint to the same block in any chain.

With this understanding of block chains, return once more to adiscussion of the location of a block with enough memory according tothe allocation algorithm of FIG. 10. In step 282 microprocessor 92determines whether every block chain has been examined for availablespace. As not a single block chain has not been examined yet,microprocessor 92 branches down to step 284 from step 282.

In step 284 microprocessor 92 determines whether every chip in thecurrent block chain has been searched for available space. If not, achip is selected from the current block chain. The first chip chosen onthe block chain is the last chip allocated into by the same process.Thus, if CPU 52 last wrote a sector into CP3, block 5, then the firstblock examined the next write will be CP3, block 5. With a blockselected, microprocessor 92 branches to step 286.

In step 286, microprocessor 92 queries the status register of thecurrent chip pair to see if the chip pair is busy. If that chip is busythen microprocessor 92 branches to step 288.

In step 288 microprocessor 92 selects the next chip on the current blockchain. For example, if the chip pair 3, block 5 was selected in step 284then chip pair 4, block 5 will be selected. This done microprocessor 92branches back up to step 284.

In step 284, microprocessor 92 again checks to see if every block on thecurrent block chain has been searched. As this is the second block onthe chain to be examined, microprocessor 92 advances again to step 286.

Here microprocessor 92 determines whether the current chip pair is busy.If so, microprocessor 92 will branch through steps 288 and 284 as justdescribed. On the other hand, if the current chip pair is not busymicroprocessor 92 proceeds to step 289.

Microprocessor 92 determines whether the currently selected chip pair isallocated to another caller. Microprocessor 92 does this in step 289 bycomparing chip and block stored in RAM for other callers to thecurrently selected chip and block. For example, let us assume that thecaller is CPU 52 and that chip pair 4, block 4 is allocated to it. If ablock in chip pair 4 is allocated to clean-up state machine 1, thenmicroprocessor 92 will not allocate memory space within chip pair 4 toCPU 52. In this situation, microprocessor 92 proceeds to step 290.

Another block on the current block chain is selected in step 290. Fromstep 292 microprocessor 92 branches back up to step 284.

As we have assumed that earlier that memory will be allocated,microprocessor 92 will eventually reach step 291 by branching throughsteps 284, 286, and 289. In step 291, microprocessor 92 determineswhether the current block selection is part of a chip that has anotherblock targeted for clean-up. If so, the current block selection isundesirable unless erasure can be suspended to program. Otherwise,writing the sector would be delayed by erasure. To avoid this,microprocessor 92 selects another chip in the current block chain instep 292. The appropriateness of that block is then determined bybranching through steps 284-291 again.

Microprocessor 92 reaches step 293 when a block that is not involved inclean-up is located. Here microprocessor 92 determines whether anotherheader can be written into BSTT 84 of the current block consistent withthe seek scheme. For example, in one embodiment only 2047 headers may bewritten into any BSTT given a maximum examination of eight headers perseek and the use of only 256 header pointers. If the maximum number ofheaders has already been written into BSTT 84, another block must beselected. Microprocessor 92 does so, by branching to step 294.Otherwise, microprocessor 92 continues its evaluation of the currentblock by branching to step 295.

In step 295 microprocessor 92 determines whether there is enough freememory within data space 86 of the current block to store the sector.This check is necessary because compression causes sector size to vary.As a result, the number of sectors that can be stored in data space 86varies. Microprocessor 92 must select another block if data space 86does not include sufficient memory to store the sector to be written.Microprocessor 92 does so in step 296.

Once a block with sufficient room to store the sector is located,microprocessor 92 performs one last check before allocating the sectordata into the current block. In step 297, microprocessor 92 examines theblock's block attribute data 90 to determine if the block is good. If itis not, microprocessor 92 advances to step 298 to select another chip onthe current block chain. On the other hand, if the block is good, anappropriate block has finally been found. Microprocessor 92 responds bybranching to step 299.

In step 299, for future reference microprocessor 92 stores in RAM, thechip and block just selected.

By branching to step 300, microprocessor 92 finally allocates memory forthe sector to be written. Microprocessor 92 does this by writing theblock sector offset into BSTT 84 of the selected block.

Microprocessor 92 updates FLASH array database 93 in step 301. The totalamount of free memory within FLASH array 62, the chip and the block(TFFA, TFFC(C_(M)) and TFFB(B_(N))) are decreased by size of the sector.When the database update is complete, microprocessor 92 branches to step302.

Finally, in step 302 the chip, block and offset to the header arereturned to the caller. As used herein, "offset to the header" or"header offset" refers to the offset from the top of the block to thetop of the header. Allocation complete, control returns to the caller instep 304.

There are a number of situations under which memory space cannot beallocated using the algorithm of FIG. 10. Perhaps memory cannot beallocated because solid state disk 60 is in standby mode or because anappropriate block cannot be located. However, and wherever, the decisionis made that memory cannot be allocated, microprocessor 92 ends up instep 306.

In step 306, microprocessor 92 determines whether its inability toallocate memory can be remedied. That depends upon the caller thatrequested memory space. If the caller is CPU 52, microprocessor 92advances to step 306. On the other hand, microprocessor branches to step310 if the caller requesting memory space is one of the two clean-upstate machines.

If the caller is CPU 52, the inability to allocate memory can beremedied by foreground clean-up. Accordingly, in step 308,microprocessor 92 initiates a foreground clean-up. As a result, CPU 52will have to wait a relatively long time before any of its writecommands are obeyed.

Very little need be done if memory cannot be allocated and the caller isone of the clean-up state machines. As will be described in detaillater, foreground clean-up uses more than one clean-up state machines.In one embodiment, foreground clean-up uses two clean-up state machinesThe first clean-up state machine runs uninterrupted until it beginserasing a block. At that time, the second clean-up state machine isstarted. It may not be possible, however, to allocate memory to thesecond clean-up state machine because so many blocks are unavailablebecause of erasure by the first clean-up state machine. This is notcatastrophic because completion of clean-up by the first clean-up statemachine will produce large amounts of free memory. Microprocessor 92prepares to await that event in step 310 by changing to idle, the statusof the calling clean-up state machine. Microprocessor 92 then returnscontrol to the caller in step 304.

FIG. 12 illustrates in flow diagram form an alternate method ofallocating free physical memory for host writes. Like the method of FIG.10, choosing a block to allocate the write into may involve up to fourdecisions. Two of the decisions for allocating memory are made inessentially the same manner as they are made according to the method ofFIG. 10. Those decisions are:

1. Is there enough reserve memory to permit the write?

2. Is this block an appropriate block to allocate the write into?

The method of FIG. 12 increases data coherency for host writes bychoosing a new block to allocate host write which has the largest amountof free memory rather than choosing the first block with enough memoryto store the sector of data to be written, as done with the method ofFIG. 10. As a result of choosing the block with the greatest amount offree memory, larger amounts of host data reside within the same blockbecause writes are allocated into the chosen block until there isinsufficient free memory to store another sector of data. This method ofallocation increases data coherency.

Allocation of free physical memory begins in step 1202 by determiningwhether minimum memory reserves within flash array 62 would bemaintained after writing the sector requested. If not, host CPU 52 isinformed that solid state disk 60 is in FLASH STANDBY in step 1204. Nofurther host writes will be permitted until microprocessor 92 releasessufficient memory to reach minimum memory reserves. On the other hand,if minimum memory reserves are available, microprocessor 92 proceeds tostep 1205.

To increase data coherency and decrease the possibility of datafragmentation, in step 1205 microprocessor 92 attempts to allocate thesector into the block host data was last written into. For simplicity'ssake, let us refer to this block as "the previous block." Microprocessor92 first determines which block is the previous block and ascertains theamount of free memory remaining in the previous block by referring todatabase 93. If the amount of free memory in the previous block is equalto or greater than the sector size, the sector will be written into theprevious block. In that case, microprocessor 92 branches ahead to step1262 to reserve memory in the previous block. On the other hand, ifthere is insufficient free memory in the previous block, microprocessor92 branches to step 1206.

In steps 1206, 1208, and 1210 microprocessor 92 initializes variablesused in the selection of a block. During these steps Most Free Block,Most Free Chip, Greatest Score and Most Free. Most Free are initializedrepresents the amount of free flash memory in the Most Free Block.Afterward, microprocessor 92 advances to step 1212.

Step 1212 begins a block by block evaluation of each block in FLASH army62. Each block is first examined to determine its appropriateness insteps 1214-1230. Once a block is found to be appropriate, thedesirability of allocating the sector write into that block isdetermined by generating a score for the block in steps 1232-1240. Thescore for each block is then compared to Greatest Score. If that scoreexceeds Greatest Score then Most Free, Most Free Block, Most Free Block,and Greatest Score are all revised to reflect the values of the currentblock in steps 1242-1248.

Choosing a block begins in step 1212 by determining whether all blockshave been examined, and thus, whether a block has been selected forallocation. Microprocessor 92 does this by determining whether all chippairs have been searched. The first pass through step 1212,microprocessor 92 selects a block for examination and branches to step1214 because not a single chip pair has been examined.

In step 1214, microprocessor 92 queries the status register of currentchip pair including the current to see if the chip pair is busy. If thatchip pair is busy, it will be unavailable for a relatively long periodof time. For this reason, microprocessor 92 branches to step 1212 toselect another block when the current block is part of a busy chip pair.Otherwise, microprocessor 92 advances to step 1216.

In step 1216, microprocessor 92 determines whether the current block ispart of a chip pair that includes another block targeted for clean-up.If so, the current block selection is undesirable unless erasure can besuspended to program. Otherwise, writing the sector could be delayed byerasure. To avoid this, microprocessor 92 returns to step 1212 toselects another block. When a current block is selected that is notincluded in a chip pair targeted for clean-up, microprocessor 92advances to step 1218.

During step 1218 microprocessor 92 determines whether memory in thecurrently selected block chip pair has been recently allocated to aclean-up state machine. Such a block is an inappropriate choice for ahost write because it increases the possibility of data fragmentation.Microprocessor 92 makes this decision by comparing chip and block storedin RAM for clean-up state machine to the currently selected chip andblock. For example, let us assume that the current block selection ischip pair 4, block 4. If a block in chip pair 4 is allocated to clean-upstate machine 1, then microprocessor 92 will not allocate memory spacewithin chip pair 4 to CPU 52. In these situations, microprocessor 92returns to step 1212 to select another block.

In step 1220 microprocessor 92 examines the block attribute data 90 forthe current block to determine if the block is good. If it is not,microprocessor 92 advances to step 1212 to select another block. On theother hand, if the block is good, microprocessor 92 responds bycontinuing its evaluation of the appropriateness of the current block.

Step 1230 is the final step in determining whether the current block isan appropriate block by assessing whether space remains in the BSTT 84of the current block. If the maximum number of headers has already beenwritten into BSTT 84, another block must be selected.

Exiting step 1230 to step 1232 means that the current block isappropriate for allocation for host writes. Thus, microprocessor 92begins evaluating the desirability of the current block, as reflected bya total score encompassing a number of factors. The first factorconsidered is the total number of free words within the current block,TFFB_(C). TFFB_(C) is multiplied by a first weight, the value of whichshould be substantially larger than the value of other weights. Thisskews the selection of a block toward the block with the greatest amountof free physical memory.

Microprocessor 92 continues evaluation of the desirability by branchingto step 1234. Microprocessor 92 examines the likelihood that the currentblock will be impacted by a future clean-up by subtracting the amount ofdirty memory within the current chip pair, designated TDFC, from thetotal possible amount of dirty memory with the current chip pair, TPDFC.The remainder of this operation is then multiplied weight 2 to generateRule 2. The smaller TDFC, the larger Rule 2 and the less likely it isthe current chip pair will soon be selected for clean-up.

The desirability of the current block is also evaluated in terms of theamount of dirty memory within the current block, TDFB. In step 1235,TDFB is multiplied by weight 3 to generate Rule 3.

Another factor evaluated in assessing the desirability of the currentblock is the number cycle counts of the current block compared to themaximum cycle count within array 62. This factor, represented by "Δcycle count" is multiplied by weight 4 in step 1236. This factor skewsselection of a block toward blocks with lower cycle counts and resultsin passive wear leveling.

Microprocessor assesses the fifth and final factor of the currentblock's desirability in step 1238. That factor is whether any otherblock in the current chip has not been allocated into. This alsoprovides some wear leveling by spreading sector data across chipboundaries. This factor is multiplied weight 5 to generate Rule 5.

Microprocessor generates a total score for the current block in step1240 by adding together Rule 1, Rule 2, Rule 3, Rule 4, and Rule 5.Microprocessor 92 then compares this total to the Greatest Score. If thetotal exceeds the Greatest Score then the current block is the mostdesirable block of the blocks that have been evaluated so far. Thus, insteps 1244, 1246, and 1248 are revised to reflect the values of thecurrent block. Afterward, microprocessor 92 returns to step 1212 toevaluate another block in FLASH array 62.

After all blocks have been evaluated, microprocessor 92 branches to step1250. There it is determined whether there is sufficient free memory inthe most free block to store the sector of data that CPU 52 wants towrite. If there is, allocation proceeds by branching to step 1262.Otherwise, microprocessor 92 branches to step 1260.

Entry to step 1260 means that FLASH array 62 must be cleaned-up before ablock with sufficient free memory to store the current sector can belocated. Thus, microprocessor 92 initiates foreground clean-up. Once ablock has been cleaned-up, microprocessor 92 returns to step 1212 toselect another block for allocation.

Eventually, a block with enough space to store the sector of data willbe identified and microprocessor 92 will reach step 1262. Allocation ofmemory space for the sector of data begins in earnest by identifying thenext available header in the BSTT 84 of the Most Free Block. Theappropriate block sector offset for the sector is then written into thatheader in step 1264. Afterward, the total amount of free flash in thearray, in the Most Free Chip and the Most Free Block are decreased bythe size of the sector. Finally, in step 1266 microprocessor indicatesthe chip, block and offset to the header now associated with the sectorof data to be written.

D. Marking Dirty a Version of a Sector

FIG. 13 illustrates in flow diagram form an algorithm for marking asector dirty, which is implemented by sectors 162. This algorithm isuseful during writes, clean-up, and creation of sector headertranslation table 94.

The mark dirty algorithm requires three pieces of information: the chip,block, and offset to the header for the sector data to be marked dirty.Given this information, the first task of microprocessor 92 in step 320is to determine the availability of the chip pair storing that versionof the sector data. Therefore, in step 320, microprocessor 92 reads thestatus register associated with the appropriate chip pair. If the chippair is busy microprocessor 92 branches to step 322. On the other hand,if the chip pair is not busy, microprocessor 92 branches to step 324.

Entry into step 322 means that the desired sector is unavailable.Clean-up is the sole possible explanation for unavailability of the chippair. Rather than waiting for clean-up to finish, microprocessor 92figuratively leaves itself a note to mark the previous copy of thesector dirty. Microprocessor 92 does this by caching the chip, block,and offset to the header of the sector data to be marked dirty. Whenclean-up is complete, the cache will be checked and the sector markeddirty. Afterward, microprocessor 92 branches to step 326 to returncontrol to its caller.

Microprocessor 92 proceeds to step 324 from step 320 if the chip pairstoring the previous copy of the sector data is not busy. Theremicroprocessor 92 marks dirty the copy of the sector pointed to by chipblock and offset. The sector of data is marked dirty by setting to 0,dirty bits in the attribute word of the sector's header. This done,microprocessor 92 proceeds to step 326, returning control to the caller.

III. Clean-up of the Solid State Disk

Solid state disk drive 60 achieves write speeds close to conventionalmagnetic disk drives by writing a sector of data to a new location eachtime it is revised, rather erasing the previous location and writing therevised data to that same physical location. As a result of thispractice, solid state disk 60 becomes sprinkled with dirty sectors.Recovering the memory space occupied by dirty sectors mandates clean-up.Stated slightly differently, the write practices of solid state diskcontroller 64 require that dirty sectors be converted into free memory.

Briefly described, clean-up involves three major tasks. First, a blockis selected as the focus of clean-up. Second, sectors of valid user dataare copied from the focus block into other blocks, referred to asdestination blocks. Third, after all valid sectors have been copied outof it, the focus block is erased, converting dirty sectors into freememory.

Clean-up is a background process, generally. Clean-up is able to run asa background process because reads and writes to FLASH array 62generally occur in bursts. As a result, there are relatively longperiods of time when microprocessor 92 is free to perform clean-up.

Clean-up is managed by a finite state machine, called a clean-up statemachine. That is, clean-up is achieved using a finite number of states,or algorithms, which are chained together. Each clean-up state points tothe next state, in effect chaining each state to another state. Eachstate takes no more than 500 μseconds of microprocessor 92 time toexecute.

Background clean-up uses a single clean-up state machine, and is grantedexecution time only when host interface 152 is inactive. In contrast,two clean-up state machines run simultaneously during foregroundclean-up of solid state disk 60. Clean-up becomes a foreground task whenwrite commands from CPU 52 cannot be executed because of lack of freememory for write requests.

Background and foreground clean-up are both handled by clean-up object164 of solid state disk controller 60. FIG. 14 illustrates in blockdiagram form the high level algorithms of clean-up object 164.Algorithms include enable CSM 340, force clean-up 342, forcemulticlean-up 344, execute one state 346, and multiCSM execute untilcondition 348.

Background clean-up is initiated when enable CSM algorithm 340 iscalled. Enable CSM 340 is called whenever the host writes to solid statememory disk 60. For example, enable CSM 340 is called in step 262 of thewrite algorithm of FIG. 9. Enable CSM 340 activates a clean-up statemachine by setting a CSM next state pointer to a clean-up state.

Execution of that first clean-up state occurs whenever top levelscheduler 150 allocates microprocessor 92 time to clean-up. Top levelscheduler 150 does so by calling execute one state algorithm 346.Execute one state 346 calls the state pointed to by the CSM next statepointer. That state, whichever it is, modifies the next state pointerprior to surrendering control of CPU 92. Thus, each time top levelscheduler 150 allocates execution time to clean-up another step will beexecuted by calling execute one state.

Foreground clean-up begins by calling service 348 of clean-up 164.MultiCSM execute until condition 348 is called when no free memory canbe allocated for a write request. This represents a decision that freememory needs to be generated before anymore writes will be executed.MultiCSM execute until condition 348 does this by enabling two clean-upstate machines.

In contrast to foreground and background clean-up, both ForceCleanUp 342and ForceMultiCleanup 344 are called by CPU 52. Force clean-up 342enables one clean-up state machine, while ForceMultiCleanup 344 enablestwo clean-up state machines.

A. Background Clean-Up 1. Initiating Background Clean-Up and the Linkingof Clean-up States

Background clean-up is initiated by a call to enable CSM 340. A flowdiagram for enable CSM 340 is shown in FIG. 15.

The first task of microprocessor 92 in step 360 is to determine whetherclean-up state machine 1 is already active. If so, nothing is to bedone. In that case, microprocessor 92 branches to step 361 without doinganything. If clean-up state machine 1 is not active microprocessor 92branches to step 362.

In step 362, microprocessor 92 performs the first of the two small tasksnecessary to begin background clean-up. There microprocessor 92 bringsclean-up state machine 1 (CSM1) active by setting a variable,CSM1.status, to active. This done, microprocessor 92 proceeds to step364.

Next, microprocessor 92 performs the second task required to set-upbackground clean-up. This is done by setting CSM1 next state pointer toEvaluate If Cleanup Is Needed state 380. The next time top levelscheduler 150 calls execute one state 346 background clean-up willbegin. Its job done, microprocessor 92 branches to step 361.

FIG. 16 is a flow diagram of execute one state 346, the algorithm thatchains the states of a clean-up state machine together. The mechanismthat allows a single algorithm to chain many states together is the CSMnext state pointer. Each clean-up state machine has its own next statepointer, which is updated at the end of each state.

Upon entry to execute one state 346, microprocessor 92 determineswhether the second clean-up state machine, CSM2, is active by examiningCSM2.status. If CSM2 is not active, microprocessor 92 branches directlyto step 370. On the other hand, CSM2 may be active as a result of aforeground clean-up that is just finishing-up. In this case,microprocessor 92 executes another state of CSM2 in step 368. Control isthen returned to the caller in step 374.

Once CSM2 is idle, microprocessor 92 will devote its attention to CSM1if CSM1 is active. Microprocessor 92 makes this determination in step370. If CSM1 is inactive, nothing is to be done but to return control tothe caller in step 374. Otherwise, the state pointed to by the CSM1 nextstate pointer is executed. Afterward, control is returned to the callerin step 374.

2. Overview of the States of a Clean-up State Machine

What is the chain of events during clean-up? Briefly described, clean-upinvolves three major tasks regardless of the type of clean-up. First, ablock is selected as the focus of clean-up. Second, on a goodsector-by-good sector basis, user data is relocated from the focus blockinto destination blocks. Relocation of user data is itself a multistepprocess, requiring allocation of memory, copying of the sector into thedestination blocks, and updating of FLASH database 93. Third, after alluser data has been copied out of it, the focus block is erased,converting dirty sectors into free memory.

FIG. 17 gives an overview of clean-up by illustrating each state of aclean-up state machine. Each bubble in FIG. 17 represents one state, oralgorithm, of a clean-up state machine. The arrows between statesrepresent the next state pointed to by the CSM next state pointer at theend of a state.

Background clean-up begins in state 380 by evaluating whether clean-upis necessary. Evaluation of whether clean-up is necessary is skippedduring foreground clean-up and forced clean-up. If clean-up is notnecessary microprocessor 92 branches down to state 396. This returnscontrol of microprocessor 92 to top level scheduler 150. On the otherhand, if clean-up is deemed necessary, a number of blocks will becleaned up. The selected number of blocks is indicated by a counter.Afterward, microprocessor 92 is pointed to state 382 by the CSM nextstate pointer.

Clean-up begins in earnest in state 382 by selecting a focus block toclean-up. The next state pointer then points microprocessor 92 to state383.

Prior to cleaning up dirty sectors within the focus block, valid sectorsof user data must be safely relocated. This task requires branchingthrough states 383, 384, 385, 386, and 388 repeatedly until each andevery sector of user data within the focus block has been safely copiedinto new locations. Relocation of user data begins in state 384 bylocating a new physical location for one good sector of user data. Theblock selected as the new location is referred to as a destinationblock. According to this terminology, clean-up has one focus block butmay have many destination blocks. Microprocessor 92 is then pointed tostate 385 by the CSM next state pointer.

In state 385, the sector is copied from the focus block into sectorbuffer 98.

In state 386 microprocessor 92 copies part of a valid sector from thesector buffer to the current destination block. Only a part of the validsector is copied at one time given write speeds and the desire to keepbackground clean-up from impacting response to read and write commandsfrom CPU 52. Thus, microprocessor 92 may pass though state 386 severaltimes before a valid sector is completely copied into destination block.Once the valid sector has been completely copied, the CSM next statepointer directs microprocessor 92 to state 388.

During state 388 microprocessor 92 updates sector header translationtable 94 so that it points to the new location of the sector just copiedfrom the focus block if the previous version of the data for the sectornumber was not marked dirty. Otherwise, microprocessor 92 marks dirtythe version of the sector it has just copied. Finally, microprocessor 92finishes writing the header associated with the new version of thesector data for the sector number. Microprocessor 92 then returns tostate 383.

Upon reentry to state 383, microprocessor 92 determines whether everygood sector within the focus block has been relocated. If not, anothergood sector will be selected for relocation, and microprocessor 92 willbranch through states 384, 385, 386, 388 and 383 until every good sectorwithin the focus block has been relocated. When that occurs, the CSMnext state pointer directs microprocessor 92 to state 390.

Microprocessor 92 begins erasure of the focus block in state 390.Microprocessor 92 initiates erasure by giving an erase command to thechip pair and indicating the block to be erased. This done,microprocessor 92 proceeds to state 392 to wait for the completion oferasure. The CSM remains in state 392 until the chip pair including thefocus block indicates completion of erasure.

The focus block erased, microprocessor 92 updates and copies blockattribute data 90 back into the focus block. This done microprocessor 92is redirected to state 380 by the CSM next state pointer.

Upon reentry to state 380, microprocessor 92 examines the block counterto determine if another focus block should be selected. If so,microprocessor 92 will branch through states 382, 383, 384, 385, 386,388, 390, 392, and 394 as described. Otherwise, microprocessor 92branches to step 396, clean-up complete.

3. Evaluating If Clean-up Is Necessary

FIG. 18 illustrates in flow diagram form the algorithm used in state 380to evaluate whether FLASH array 62 should be cleaned-up. Using thisalgorithm, background clean-up is triggered when total dirty memoryrises above a set percentage of total memory reserves. Total memoryreserves are defined as the sum of total free memory (TFFA) and totaldirty memory (TDFA). The percentage of total dirty to total reservesaffects the performance and power efficiency of solid state disk 60.

Microprocessor 92 begins in step 400 by determining whether evaluationis necessary. This is done by examining a counter to determine thenumber of blocks remaining to clean-up. Any number greater than zeroindicates that clean-up has already been triggered and furtherevaluation is not necessary. The first time microprocessor 92 entersstate 380 the counter will be less than zero, forcing microprocessor 92to branch to step 402.

In step 402 microprocessor 92 calculates the total amount of dirtymemory within FLASH array 62 as a percentage of total memory reserves.

In step 404 the percentage calculated in step 402 is compared to aclean-up trigger point. The clean-up trigger point may be fixed at acertain percentage of total memory reserves or it may be adaptivelyvaried based upon the demands of CPU 52. An algorithm for adaptivelyvarying the clean-up trigger point will be described in detail later. Ifthe total dirty is equal to, or exceeds, the trigger pointmicroprocessor 92 sets background clean-up in motion by branching tostep 406. On the other hand, if the percentage of total dirty memorydoes not exceed the trigger point microprocessor 92 branches to step409. Because clean-up is not necessary, the status of CSM1 is set toidle. Control is returned to top level scheduler 150 without initiatingclean-up by branching to state 410.

Microprocessor 92 sets the block counter in step 406 to the maximumnumber of blocks to be cleaned-up. In one embodiment, that maximum isset to three. This done, microprocessor 92 advances to step 408.

In step 408, microprocessor 92 advances the CSM next state pointer tothe next state in the clean-up process, state 382. This done,microprocessor 92 returns control to top level scheduler 150 bybranching to step 410.

During subsequent calls to evaluate if clean-up is necessary, state 380microprocessor 92 begins by examining the block counter in step 400. Ifthe block counter exceeds zero the need for clean-up will not beevaluated. Instead, microprocessor 92 decrements the block counter andbranches to step 401.

In step 401 microprocessor 92 points the CSM next state pointer to state382, choose block to clean up. Microprocessor 92 then exits state 380via return step 410.

FIG. 19 illustrates in flow diagram form an algorithm for adaptivelyvarying the trigger point used in step 404 to evaluate whether clean-upis necessary. The algorithm shown belongs to sectors 162.

In step 412, microprocessor 92 determines whether the trigger point hasbeen set too low; that is, whether FLASH array 62 is being allowed toget too dirty. Microprocessor 92 does this by determining whether somenumber P of foreground clean-ups have occurred in the last M writes,where M is a large number. If so, microprocessor 92 advances to step414. Otherwise, microprocessor 92 branches to step 416.

In step 414 microprocessor 92 decreases the trigger point by a smallamount, because FLASH array 62 has been allowed to get too dirty. Theeffect of this is to keep FLASH array 62 from becoming as dirty in thefuture. From step 414 microprocessor 92 returns control to the callingprogram.

In step 416, microprocessor 92 considers whether FLASH array 62 is beingkept too clean. Microprocessor 92 determines this by evaluating thenumber of foreground erases in the last N writes, where N is a largenumber. If there has not been a foreground erase in a very long timemicroprocessor 92 branches to step 418. Otherwise, microprocessor 92proceeds to step 420 without changing the trigger point.

Microprocessor 92 increases the trigger point slightly in step 418. Thisallows FLASH array 62 to become slightly dirtier before backgroundclean-up is triggered, which improves clean-up efficiency. The triggerpoint adjusted, microprocessor 92 returns control to the calling programby branching to step 420.

4. Choosing a Block to Clean-Up

Once clean-up is triggered, a focus block is chosen in state 382. Thegoal in choosing a focus block is to select the block that it is mostcost effective to clean. Cost effective clean-up involves striking abalance between keeping FLASH array 62 so clean that power consumptionis excessive and keeping FLASH array 62 so dirty that foregroundclean-up frequently occurs. A number of factors are to be considered indetermining whether a block is a cost effective focus for clean-up.First, how dirty is the block? The dirtier a block is the morecost-effective it is to erase that block. In selecting a focus, thisfactor is weighted heaviest to skew the decision toward minimizingforeground clean-ups. Second, how much free memory is in the chip pairthat the focus is part of? This factor is considered only for FLASHmemory devices without the ability to suspend erasure to program. InFLASH devices without the ability to suspend erasure, every block withinthe chip pair is unavailable during the erasure of a single block. Thisreduces the desirability of cleaning up a particular block if lots offree memory resides within the same chip pair. Third, how much validuser data, or good data, is there within the focus block? This alsoreduces the desirability of choosing a particular block as focus becauseit increases the number of sectors that must be copied out. Statedconversely, the fewer good sectors that must copied out the moredesirable a block is to clean-up. Other factors may be considered whenchoosing a focus block.

FIG. 20 illustrates an algorithm for choosing a block to clean-up instate 382 which uses the factors discussed above. Briefly described, ablock is chosen as a focus of clean-up by generating a score for eachblock within FLASH array 62. The score is generated by appropriatelyweighting each of the factors and summing the weighted factors together.The block with the highest total score is selected as the focus ofclean-up.

Choosing a focus block begins in step 440 by determining whether allblocks have been examined, and thus, whether a focus block has beenselected. Microprocessor 92 does this by determining whether all chippairs have been searched. The first pass through step 440,microprocessor 92 branches to step 444 because not a single chip pairhas been examined.

In step 444 microprocessor 92 selects a block to examine within theselected chip pair. Microprocessor 92 then advances to step 445 if everyblock has not been examined yet.

In step 445, microprocessor 92 determines whether the block is part of achip pair that is currently the focus of clean-up by another clean-upstate machine. This makes the current block a less desirable focus blockbecause all blocks within a chip targeted for clean-up will beunavailable at some point for a long period of time. Microprocessor 92is able to determine whether a chip is targeted for clean-up by anotherclean-up state machine by examining the chip and block informationstored in RAM. After multiplying this factor by weight 0 to generateRule 0, microprocessor 92 then advances to step 446. The first score isgenerated by multiplying the total number of dirty words within thecurrent block TDFB_(C), by a first weight. The value of this weightshould be substantially larger than the value of subsequent weights.Microprocessor 92 then branches to step 448 from step 446.

Microprocessor 92 calculates a second score for the current block instep 448. There the total amount of free FLASH within the selected chippair TFFC_(C) is multiplied by a second weight. Afterward,microprocessor 92 branches to step 449.

A third score is calculated in step 449. The desirability of a block isevaluated there in terms of whether the current block resides within thechip pair last erased. Afterward, microprocessor 92 advances to step450.

In step 450 microprocessor 92 calculates a fourth score for the currentblock. Microprocessor 92 begins calculating this score by subtractingthe total number of good bytes in the current block from the totalnumber of bytes within that block. Microprocessor 92 then multipliesthis number by a fourth weight. The fourth weight is meant to biasslightly the selection of a focus block towards choosing a block withless good sector data. Thus, microprocessor 92 then branches to step451. The fourth weight is relatively small compared to the first weight.

The fifth and final score for the current block is generated during step451. The fifth factor considered skews selection toward blocks withlower cycle counts, thus, performing passive wear leveling. The fewercycle counts of the current block compared to the maximum cycle countwithin array 62, the more desirable the block is. This factor,represented by "Δ cycle count" is then multiplied by the fifth weight.

In step 452 microprocessor 92 generates a total score for the block byadding together Rules 1, 4, and 5, which reflect desirablecharacteristics of the current block. From this sum Rules 0, 2, and 3are subtracted because they indicate undesirable characteristics of afocus block.

Advancing to step 454, microprocessor 92 then compares the total scorefor the current block to the greatest total score. If the total scorefor the current block exceeds the greatest total score, then thegreatest total score is revised upward and microprocessor 92 advances tostep to 456. On the other hand, microprocessor 92 branches back up tostep 444 if the total score for the current block is not greater thanthe greatest total score.

In step 456 microprocessor 92 stores the chip, and block and score ofthe block with the current greatest total score. Microprocessor 92 thenreturns to step 444.

Upon reentry to step 444 microprocessor decides whether every blockwithin the currently selected chip pair has been evaluated. If not,microprocessor 92 branches through steps 445, 446, 448, 449, 450, 451,452, 454, 456, and 444 until every block within the selected chip pairhas been evaluated. Microprocessor 92 branches back to step 440 onceevery block within the selected chip pair has been evaluated.

Step 440 determines whether the blocks within every chip pair have beenexamined. If not, microprocessor 92 branches through steps 444, 445,446, 448, 449, 450, 451, 452, 454, 456, and 440 until all blocks in allchip pairs within FLASH array 62 have been evaluated. Microprocessor 92then proceeds to step 458.

Microprocessor 92 writes the chip and block selected as the focus ofclean-up into the focus window. Microprocessor 92 uses the block andchip address stored in step 456 to do this. Afterward, microprocessorproceeds to step 460.

Having reached its current goal, in step 460 microprocessor 92 sets theCSM next state pointer to allocate free physical sector, state 384.

From step 460 microprocessor 92 returns control to top level scheduler150 by branching to step 462.

5. Allocating Free Physical Sectors for Clean-up

Prior to erasing the focus block, all valid sectors of user data must becopied from the focus block and relocated into destination blocks. Thefirst step in the relocation process is allocating memory space for thesectors of the user data currently residing in the focus block. FIG. 21illustrates an algorithm for state 384, which manages the allocation ofmemory during clean-up.

Microprocessor 92 begins allocation in step 472 by determining whetherevery good sector within the focus block has been relocated.Microprocessor 92 makes its decision by scanning down BSTT 84 of thefocus block until reaching the next header. Every sector of user datahas been copied out of the focus block if the attribute word in theheader is stack 88 which separates the BSTT 84 and the data space 86 andis represented by FFFF (hexadecimal). In that case, microprocessor 92branches to state 480. If every valid sector has not been copied out,microprocessor 92 examines the attribute word to see if the sector isvalid or dirty. Dirty sectors cause microprocessor 92 to scan down BSTT84 to the next header. Microprocessor 92 responds to good sectors bydetermining the sector size and advancing to step 474.

In step 474 microprocessor 92 allocates memory for the sector identifiedin step 472 by calling the allocate algorithm of FIG. 10 or an alternatealgorithm to be discussed later. Afterward, microprocessor 92 advancesto step 476.

Microprocessor 92 prepares to exit state 384 in step 476 by resettingthe CSM next state pointer to state 385. Thus, when background clean-upresumes execution, copying of the sector from the focus block intosector buffer 98 will begin. This done, microprocessor 92 returnscontrol to top level scheduler 150 by branching to step 478.

Microprocessor 92 returns again and again to state 476 until every validsector has been copied out of the focus block. Once that occurs,microprocessor 92 branches to step 480 from step 472. In step 480microprocessor 92 initiates focus block erasure by appropriately settingthe CSM next state pointer to start erase. Control is returned to thecalling program.

6. Allocating Free Physical Memory for Clean-up

FIG. 22 illustrates in flow diagram form an alternate method ofallocating free physical memory for clean-up. A block selected as adestination block for clean-up must be both appropriate and desirable.The appropriateness of a block is assessed in essentially the samemanner discussed previously with respect to the allocation methodsillustrated in FIGS. 10 and 12. However, the assessment and definitionof a desirable destination block differs. The most desirable destinationblock is the block that most closely fits the sectors of user data to becopied out of focus block. Defining the desirability of a destinationblock in this manner increases the coherency of user data by preservinglarge amounts of free data for host writes.

Allocation of memory during clean-up begins with initialization ofvariables used in the selection of the destination block. In steps 2202and 2204 Best Fit Block and Best Fit Chip are both set to 0.

In step 2206 microprocessor 92 calculates the amount of free memorynecessary to store the remaining sectors of user data within the focusblock. Thus, even though memory for only one sector is allocated at atime, desirability is evaluated in terms of a block capable of storingall the remaining sectors of user data within the focus block. (As usedherein, "sectors of user data" and "user data" is defined as memory thatis neither free, nor dirty.)

In step 2207 microprocessor 92 initializes one more variable prior tobeginning its evaluation of possible destination blocks. Score isinitialized to a maximum value representative of the maximum amount offree memory in any block. Score is revised during evaluation to indicatethe score of the most desirable destination block; that is the blockwhose available free memory is closest to the total free memory needed.This done, microprocessor 92 advances to step 2208.

The appropriateness of each block as a destination block is evaluated insteps 2208-2216 in the same manner discussed with respect to FIG. 12.Therefore, these steps are not discussed in detail.

Eventually an appropriate block will be identified and microprocessor 92will reach step 2218. There the desirability of the current block as adestination block will be determined. One factor defines destinationblock desirability according to the service of FIG. 22. How closely doesthe amount of free memory in the current match the amount of free memoryneeded to store all sectors of user data remaining within the focusblock? A variable, Total, measures the desirability of the current blockin terms of the free memory remaining after storing the user data withinthe focus block. Total may be a positive number or a negative number,depending upon whether the free memory in the current block is to greator too little. Microprocessor 92 then branches to step 2220.

In step 2220 microprocessor 92 determines whether the current block is abetter fit than the block currently selected as the best fit. In otherwords, the absolute value of Total_(CurrentBlock) is compared to theabsolute value of Score. Using absolute values in the comparison ensuresthat blocks that are too small are considered as well as blocks thathave more free memory than is strictly necessary. Microprocessor 92returns to step 2208 to evaluate another block if the absolute value ofTotal_(CurrentBlock) is not less than the absolute value of Score. Onthe other hand, if the current block is a better fit than a previousblock then microprocessor 92 branches to step 2222 to revise best fitvalues.

Score is reset to Total_(CurrentBlock) in step 2222. Afterward, in steps2224 and 2226 Best Fit Chip and Best it Block are revised so that theypoint to the current block. This done, microprocessor 92 returns to step2208 to evaluate another block.

After evaluating each block within FLASH array 62 microprocessor 92exits from step 2208 to step 2238, microprocessor 92 then determineswhether there is sufficient memory in Best Fit Block to store the sectorof data currently being moved out of the focus block. If not,microprocessor 92 idles the current clean-up machine in step 2230.Otherwise, microprocessor 92 advances to step 2232, 2234 and 2236 tobegin the process of actually reserving memory in the Best Fit Block,which it does in the same fashion as discussed previously with respectto FIGS. 10 and 12.

7. Copying Good Sectors

Copying a sector from the focus block into a destination block is a twostep process. The selected sector of valid user data is first copiedfrom the focus block into sector buffer 98. Secondly, the valid sectoris copied from sector buffer 98 into the destination block. This twostep process is not necessary, but it improves the reliability of userdata by taking advantage of the error detection capabilities of an errordetection circuit that is in the path of sequencer 168.

The algorithm of FIG. 23 handles the copying of a sector of user datafrom the focus block to sector buffer 98. First, in step 480, the sectoris moved into sector buffer 98. Second, in step 482, microprocessor 92sets the next state pointer to copy par of sector. Control is thenreturned to the caller in step 484.

Copying a sector from its temporary refuge in sector buffer 98 into thedestination block is the responsibility of state 386, called Copy Partof Sector. Only part of the sector is copied each time state 386 isexecuted to ensure that its execution time is less than a selectedmaximum execution time. In one embodiment the maximum execution time ofeach state is limited to 500 μseconds.

FIG. 24 illustrates an algorithm for copying part of a sector fromsector buffer 98. Microprocessor 92 begins in step 490 by deciding ifthe rest of the sector can be copied into the destination block withinthe maximum execution time. In one embodiment, microprocessor 92 doesthis by comparing to 20 the number of bytes remaining to be copied fromsector buffer 98. The maximum number of words that can be copied willvary depending upon the selected maximum execution time. Microprocessor92 branches to step 492 if it cannot copy the remainder of the sectorwithin the maximum execution time.

In step 492 microprocessor 92 copies 20 bytes from sector buffer 98 tothe destination block indicated by the clean-up state machine'sdestination window. Afterward, microprocessor 92 branches to step 494.

In step 494 microprocessor 92 points the CSM next state pointer back tostate 386, guaranteeing that all of the sector will eventually be copiedout of sector buffer 98. Control is returned to top level scheduler 150when microprocessor 92 branching to step 500.

Eventually there will be a pass through state 386 in which all remainingwords can be copied into the destination block within the maximumexecution time. During this pass, microprocessor 92 branches from state490 to state 496.

Microprocessor 92 copies all the words remaining in sector buffer 98into the destination block in step 496. Afterward, microprocessor 92advances to step 498.

Having completely relocated the valid sector into its destination block,microprocessor 92 prepares to update sector header translation table 94to reflect the sector's new physical location. At this point,microprocessor 92 also finishes writing the header for the new sector ofdata. This is done in step 498 by resetting the CSM next state pointerto state 388. Copying of the sector finished, microprocessor 92 exitsstate 386 via step 500.

8. Updating Databases After Copying a Sector

Updating of BSTTs 84 and SHTT 94 is delayed until after a sector of userdata has been completely relocated from the focus block to thedestination block. This ensures that CPU 52 always has access to a validcopy of the sector, even while the sector is being relocated. Thus, CPU52 is free to re-write the sector while the sector is being moved. As aresult, there can be three or more versions of the same sector of datawithin FLASH array 62, only one version of which is valid. When thisoccurs, microprocessor 92 has effectively copied a dirty sector duringstate 386. The Post Copy Database Update algorithm of state 388anticipates this problem by determining which databases should beupdated and how they should be updated.

FIG. 25 illustrates a flow diagram for updating databases during state388. Microprocessor 92 begins this task in step 502 by determiningwhether the host has written the sector while the CSM was copying it.Microprocessor 92 can determine this by checking if the header for thesector just relocated has been marked dirty. Let us call thisheader_(FOCUS). Once microprocessor 92 knows this, it knows whichtranslation tables should be updated and how they should be updated.There are two possibilities.

Consider first the train of events when header_(FOCUS) is not markeddirty. In this situation, sector header translation table 94 and blocksector translation tables 84 for both the destination block and focusblock must be updated. To begin to do so microprocessor 92 branches fromstep 502 to step 506.

Database updates begin with the destination block. In step 506,microprocessor 92 writes the LSN into BSTT 84 of the destination block.Afterward, microprocessor 92 branches down to step 508.

Microprocessor 92 changes the address of the sector within SHTT 94 sothat it points to the header in destination block, rather than theheader in the focus block. Microprocessor 92 then advances to step 510.

Microprocessor 92 completes the updating of all three translation tablesin step 510. There microprocessor 92 marks dirty the header in the focusblock. Microprocessor 92 does so by calling the mark dirty algorithm ofFIG. 13. Database updates complete, microprocessor 92 branches to step512.

Entry to step 512 indicates copying of a sector is complete.Microprocessor 92 prepares to scan the focus block for another goodsector to copy in step 512 by setting the CSM next state pointer toallocate free physical sector state 384.

Afterward, control of microprocessor 92 is returned to top levelscheduler 150 by branching to step 514.

What happens if the sector is revised while it is being relocated to thedestination block? As it has already been marked dirty, header_(FOCUS)need not be revised. Nor does SHTT 94 require revision because italready points to the header associated with the revised sector. OnlyBSTT 84 of the destination block requires revision to reflect that theheader in the focus block is invalid. Microprocessor 92 begins this taskby branching from step 502 to step 516.

Microprocessor 92 updates BSTT 84 of the destination block in step 516by calling the mark dirty algorithm of FIG. 13. As before,microprocessor 92 returns control to top level scheduler 150 bybranching block 514.

9. Erasing the Focus Block

Conversion of dirty sectors into free memory via erasure begins aftercopying all good sectors out of the focus block. Erasure of the focusblock involves initiating erasure and then waiting for erasure tofinish.

FIG. 26 illustrates an algorithm for initiating erasure in state 390.Prior to issuing an erase command, in step 520 microprocessor 92preserves block attribute data 90 in step 520 by copying it to sectorbuffer 98.

At last, in step 522 microprocessor 92 issues the erase command to thechip pair including the focus block. Nothing remains to be done now butto wait.

Microprocessor 92 prepares for its long wait in step 524 by setting theCSM next state pointer to wait for erase complete state 392.Microprocessor 92 then returns control to the calling program bybranching to state 526.

FIG. 27 illustrates an algorithm used to wait for and detect thecompletion of erasure. In step 530, microprocessor 92 determines whethererasure is still on-going by querying the chip pair's status register.As long as the status register indicates that erasure is not complete,microprocessor 92 branches to step 532. There microprocessor 92 preparesto continue its wait by setting the CSM next state pointer to wait forerase complete state 392. Once erasure is complete, microprocessor 92branches to state 534, where the CSM next state pointer is set to posterase update state 394.

10. Posterase Database Update

A few tasks require attention after the focus block has been erased.These tasks are managed by posterase database update state 394, which isillustrated in FIG. 28.

The first task tackled by microprocessor 92 in step 536, is updating theFLASH array database to reflect the increased number of cycle countscaused by the preceding erasure.

The second task performed by microprocessor 92 is the restoration to thefocus block of its block attribute data 90 with the new cycle count,which was stored in sector buffer 98 during erasure. Microprocessor 92does this in step 538.

Microprocessor 92 performs its last update task in step 540. Theremicroprocessor 92 clears the semaphore stored in RAM to indicate thatthe block is no longer a focus of clean-up. As a result, the allocatealgorithm of FIG. 11 is free once again to allocate sectors, ifpossible, into the chip pair including the former focus block.

Before exiting state 394, in step 542, microprocessor 92 resets the CSMnext state pointer to state 380, Evaluate If Clean-Up is Necessary.

B. Foreground Clean-Up

The goal of foreground clean-up is to generate rapidly a large amount offree memory within FLASH array 62. Foreground clean-up achieves its goalby simultaneously operating two clean-up state machines and forcing CPU52 to wait for free memory. Thus, foreground clean-up differs frombackground clean-up by impacting the computer user. Foreground clean-upalso differs in the manner in which it is triggered as compared tobackground clean-up. When active, background clean-up is initiated aftereach write of a sector by CPU 52. Foreground clean-up is triggeredduring the allocation of free physical memory. If sufficient free FLASHmemory to write a sector cannot be allocated then a call to MultiCSMexecute until condition 348 begins foreground clean-up. Additionally,foreground clean-up differs from background clean-up by foregoingevaluation of whether clean-up is necessary, that decision effectivelybeing made during allocation of free memory. Foreground clean-uptherefore begins by choosing a block to clean-up.

FIG. 29 illustrates the algorithm used to manage foreground clean-up,MultiCSM Execute Until Condition 348. Using this algorithm, the firstclean-up state machine is activated and allowed to run continuouslyuntil it reaches wait for erase complete state 392. At that point, thesecond clean-up state machine is activated and allowed to runcontinuously until it reaches wait for erase complete state 392. Thefirst clean-up state machine then completes the clean-up its focus blockwithout interruption. With one block clean, and another block to beclean shortly, microprocessor 92 again responds to commands from CPU 52.

Foreground clean-up begins in step 550 where microprocessor 92 sets toone block counts for both the first clean-up state machine and thesecond clean-up state machine. Recall from the discussion of FIG. 17that the block count determines how many blocks are cleaned-up, and thushow long clean-up lasts. Setting both block counts to one minimizesforeground clean-up execution time. Microprocessor then advances to step552.

With step 552 microprocessor 92 begins the process of activating bothclean-up state machines, which continues through step 562.Microprocessor 92 determines whether the first clean-up state machine,CSM1, is active in step 552 by examining the status of CSM1. If CSM1 isnot active, microprocessor 92 branches to state 554 to activate it. Onthe other hand, if CSM1 is already active, microprocessor 92 branches tostep 558 to activate the second clean-up state machine, CSM2.

Microprocessor 92 changes the status of CSM1 from idle to active in step554. Microprocessor 92 proceeds from step 554 to step 556.

Microprocessor 92 prepares CSM1 to begin clean-up by setting its nextstate pointer to choose block to clean up state 382. This done,microprocessor 92 turns its attention to CSM2 by branching to step 558.

Microprocessor 92 determines whether CSM2, is active in step 558. If so,microprocessor 92 proceeds to step 564 to start running CSM1. Otherwise,microprocessor 92 branches to step 560.

Microprocessor 92 sets CSM2 to its active state in step 560. Afterward,in step 562, microprocessor 92 prepares CSM2 to begin clean-up bysetting its next state pointer to choose block to clean up state 382.

Set-up of the two clean-up state machines complete, microprocessor 92 isfinally able to start one up. In step 564, microprocessor 92 determineswhich clean-up state machine should be given execution time.Microprocessor 92 does this by determining whether CSM1 has begunerasing a block. If CSM1 has not, microprocessor 92 proceeds to step566.

Microprocessor 92 executes one state of CSM1 in step 566. Microprocessorexecutes CSM1 states until wait for erase complete state 392.Microprocessor 92 reaches states 392 by looping repeatedly through steps566 and 564. When CSM1 reaches wait for erase state 392, microprocessor92 turns its attention to more fruitful fields by branching to state568.

Microprocessor 92 determines whether CSM2 requires additional executiontime to finish cleaning its focus block in step 568. If so,microprocessor 92 branches to step 570 where a single state of CSM2 isexecuted. Microprocessor executes CSM2 states until wait for erasecomplete state 392. At this point, microprocessor turns its attentiononce again to CSM1 by branching from step 568 to step 572.

In step 572 it is determined whether CSM1 has completed its clean-up ofits focus block. Microprocessor 92 does this by examining the statusword associated with CSM1. If CSM1 is not idle, microprocessor 92branches to step 574 to execute another CSM1 state. By branching backand forth between steps 572 and 574, microprocessor 92 eventuallycompletes the clean-up managed by CSM1. At this point, one block hasbeen completely cleaned-up and another block will shortly be cleaned-up.With a large reserve of free memory imminent, microprocessor 92 returnscontrol to top level scheduler 150 by branching to step 576. Foregroundclean-up thus ceases.

C. Forced Clean-up

Forced clean-up is yet another way of activating and controllingclean-up state machines. Forced clean-up differs from both foregroundclean-up and background clean-up in that it is initiated by a commandfrom the computer user. In this type of clean-up, the computer userdecides that clean-up is necessary without evaluating the total amountof free memory, as done in state 380.

Solid state disk controller 64 recognizes two types of forced clean-up.One type of forced clean-up activates only one clean-up state machine.This type of clean-up is managed by a service named Force Cleanup 342.The other type of clean-up activates two clean-up state machines and ismanaged by a service called ForceMultiCleanup 344.

FIG. 30 illustrates an algorithm for Force Cleanup 342. The task ofForceCleanup is relatively simple: activate and start-up CSM1, ifnecessary.

In step 580 microprocessor 92 begins its task by examining the status ofCSM1 to see if it is active. If so, microprocessor 92 branches to step587. If CSM1 is idle, then microprocessor 92 branches to step 582.

Microprocessor 92 enables CSM1 in step 582 by changing its status toactive. Microprocessor 92 then forces clean-up to begin by bypassingevaluate if clean-up is necessary state 380 and setting CSM1 next statepointer to choose block to clean-up 382. Its responsibilities fulfilled,microprocessor 92 branches from step 584 to step 586.

Microprocessor 92 ensures that CSM1 initiates clean-up in step 587, ifit has not already done so. Microprocessor 92 does so by examiningCSM1's next state pointer and if it is equal to evaluate clean-upnecessary branching to step 584. This circumvents any evaluation thatclean-up is unnecessary.

FIG. 31 illustrates an algorithm for Force Multi Cleanup 344. The taskof this algorithm is to activate both clean-up state machines, ifnecessary.

The first two steps of Force Multi Cleanup 344 echo the steps of ForceCleanup 342. In step 590, microprocessor 92 determines whether CSM1 isactive. If it is not, in step 592 CSM1 is enabled and choose block tocopy state 382 is selected as the next state of CSM1. On the other hand,if CSM1 is already active, microprocessor 92 branches to step 591.

Microprocessor 92 determines whether CSM1 has begun clean-up in step591. Recall that clean-up of a block is not initiated each time CSM1 isactivated. Clean-up has begun if the next state pointer points to anystate other than Evaluate If Clean-Up is necessary. In this casemicroprocessor 92 branches directly to step 594. Otherwise,microprocessor 92 proceeds to step 592.

There microprocessor 92 forces CSM1 to clean-up a block by resetting theCSM1 next state pointer to Choose Block to Clean-Up. Afterward,attention turns to CSM2 by the branch to step 594 from step 592.

Microprocessor 92 enables the second clean-up state machine, CSM2, instep 594. Microprocessor 92 then bypasses evaluate if clean-up isnecessary state 380 and sets CSM2 next state pointer to choose block tocopy state 382. Microprocessor 92 branches from step 594 to step 596,returning control to top level scheduler 150.

IV. Summary

A method of locating sectors of data in a solid state memory disk hasbeen described. The sector number associated with a sector of data isused to find a pointer into a block sector translation table. Headerswithin the block sector translation table near the pointer are examineduntil a header including a logical sector number equal to the sectornumber is found, or a maximum scan count is reached, whichever occursfirst.

In the foregoing specification, the invention has been described withreference to specific exemplary embodiments thereof. It will, however,be evident that various modifications and changes may be made theretowithout departing from the broader spirit and scope of the invention asset forth in the appended claims. The specification and drawings are,accordingly, to be regarded in an illustrative rather than a restrictivesense.

What is claimed is:
 1. A method of locating a sector of data stored in asolid state memory disk including a nonvolatile semiconductor memorydevice having a block, the block including a block translation tablestoring headers and a data space storing sectors of data, the methodcomprising the steps of:a) receiving an input sector number and apointer into the block translation table; b) initializing a scan count;c) selecting a selected header from a multiplicity of headers near alocation pointed to by the pointer, each of the multiplicity of headersincluding a selected logical sector number, a selected offsetrepresenting an offset to a selected sector of data associated with theselected logical sector number; d) indicating the selected offset if theselected logical sector number of the selected header equals the inputsector number; e) incrementing the scan count if the selected logicalsector number does not equal the input sector number; f) comparing thescan count to a maximum scan count; and g) repeating steps c) through f)once for each of a set of headers which excludes the selected header ifthe scan count is not equal to the maximum scan count.
 2. The method ofclaim 1 wherein the nonvolatile semiconductor memory device is a FLASHEEPROM.
 3. A method of locating a sector of data stored in a solid statememory disk including a nonvolatile semiconductor memory device having ablock, the block including a block translation table storing headers anda data space storing sectors of data, the method comprising the stepsof:a) receiving an input sector number and a first pointer into theblock translation table; b) initializing a scan count; c) selecting aselected header from a multiplicity of headers near a location pointedto by the first pointer, each of the multiplicity of headers including aselected logical sector number, a selected offset representing an offsetto a selected sector of data associated with the selected logical sectornumber, and including a selected cyclical redundancy check for theselected header; d) if the selected logical sector number of theselected header equals the input sector number:1) generating a firstcyclical redundancy check for the selected header; 2) indicating theselected offset and ending if the selected cyclical redundancy checkequals the first cyclical redundancy check; 3) incrementing the scancount if the first cyclical redundancy check does not equal the selectedcyclical redundancy check; e) if the selected logical sector number doesnot equal the input sector number incrementing the scan count; f)comparing the scan count to a maximum scan count, the maximum scan countequal to a maximum number of headers in the block translation tabledivided by a maximum number of pointers into the block; and g) repeatingsteps c) through f) once for each of a set of headers which excludes theselected header if the scan count is not equal to the maximum scancount.
 4. The method of claim 3 wherein the nonvolatile semiconductormemory device is a FLASH EEPROM.
 5. The method of claim 3 furthercomprising the steps of:h) marking the selected header dirty if theselected logical sector number is equal to the input sector number andthe selected cyclical redundancy check does not equal the first cyclicalredundancy check.
 6. The method of claim 3 further comprising the stepof:h) indicating that a first sector of data associated with the inputsector number cannot be found if the scan count equals the maximum scancount.
 7. The method of claim 3 prior to the step of initializing thescan count further comprising the steps of:a.1) reading a next headerlocated in a predetermined location relative to a last located header inthe block translation table, the next header including a next logicalsector number, a next offset representing an offset to a next sector ofdata associated with the next sector number; a.2) indicating the nextoffset if the next logical sector number equals the input logical sectornumber; and a.3) continuing locating the sector of data if the nextlogical sector number does not equal the input logical sector number. 8.The method of claim 7 further comprising the steps of:h) comparing thefirst pointer to an invalid pointer value; and i) indicating a firstsector of data associated with input sector number cannot be found andending if the first pointer is equal to the invalid pointer value. 9.The method of claim 3 prior to the step of initializing the scan countfurther comprising the steps of:a.1) reading a next header located in apredetermined location relative to a last located header in the blocktranslation table, the next header including a next logical sectornumber, a next cyclical redundancy check for the next header, a nextoffset representing an offset to a next sector of data associated withthe next logical sector number; a.2) if the next logical sector numberequals the input sector number1) generating a second cyclical redundancycheck for the next header; 2) indicating the next offset if secondcyclical redundancy check equals the next cyclical redundancy check; 3)indicating that a first sector of data associated with the input sectornumber cannot be found if the second cyclical redundancy check does notequal the next cyclical redundancy check; and a.3) continuing locatingthe sector of date if the next logical sector number does not equal theinput sector number.
 10. A method of locating a sector of data stored ina solid state memory disk including a nonvolatile semiconductor memorydevice having a block, the block including a block translation tablestoring headers and a data space storing sectors of data, comprising thesteps of:a) receiving an input sector number and a pointer into theblock translation table; b) initializing a scan count; and c) repeatingthe steps ofc.1) selecting a selected header from a multiplicity ofheaders near a location pointed to by the pointer, each of themultiplicity of headers including a selected logical sector number, aselected offset representing an offset to a selected sector of dataassociated with the selected logical sector number; c.2) indicating theselected offset if the selected logical sector number of the selectedheader equals the input sector number; and c.3) incrementing the scancount if the selected logical sector number does not equal the inputsector number; until a maximum scan count is reached or the selectedlogical sector number of the selected header equals the input sectornumber.
 11. A method of locating a sector of data stored in a solidstate memory disk including a nonvolatile semiconductor memory devicehaving a block, the block including a block translation table storingheaders and a data space storing sectors of data, comprising the stepsof:a) receiving an input sector number and a pointer into the blocktranslation table; b) initializing a scan count; and c) repeating thesteps ofc.1) selecting a selected header from a multiplicity of headersnear a location pointed to by the pointer, each of the multiplicity ofheaders including a selected logical sector number, a selected offsetrepresenting an offset to a selected sector of data associated with theselected logical sector number; c.2) if the selected logical sectornumber does not equal the input sector number incrementing the scancount c.3) if the selected logical sector number of the selected headerequals the input sector number,c.3.a) generating a first cyclicalredundancy check for the selected header; c.3.b) indicating the selectedoffset and ending if the selected cyclical redundancy check equals thefirst cyclical redundancy check; and c.3.c) incrementing the scan countif the first cyclical redundancy check does not equal the selectedcyclical redundancy check; until a maximum scan count is reached or theselected logical sector number of the selected header equals the inputsector number and the selected cyclical redundancy check equals thefirst cyclical redundancy check.
 12. A method of locating a sector ofdata stored in a solid state memory disk including a non-volatile memorydevice which includes an organizational data structure storing aplurality of data descriptors and a data space for storing a pluralityof sectors described by the plurality of data descriptors, comprisingthe steps of:a) receiving an input sector number and a pointer to aninitial location in the organizational structure; b) initializing a scancount; c) reading a data descriptor for a stored sector, the datadescriptor containing a stored sector number and an offset of the storedsector into the data space; d) incrementing the scan count if the storedsector number does not correspond to the input sector number; e)repeating steps c) through e) for a plurality of different datadescriptors until the scan count reaches a maximum count if the storedsector number does not correspond to the input sector number; and f)indicating the offset if the stored sector number corresponds to theinput sector number.